Hardware transaction transient conflict resolution

ABSTRACT

In an approach for resolving terminated transactions in a transactional memory environment, a processor initiates a hardware transaction in a computing environment, wherein the hardware transaction accesses a memory location, and wherein the hardware transaction includes a transaction begin indicator and a transaction end indicator. A processor detects a conflicting access of the memory location while executing the hardware transaction. A processor aborts the hardware transaction based on the conflicting access of the memory location. Hardware determines that the conflicting access of the memory location is a transient condition. A processor reinitiates the hardware transaction.

BACKGROUND

This disclosure relates generally to transactional memory systems andmore specifically to a method, computer program and computer system forimproving the efficiency of transactional instruction processing.

The number of central processing unit (CPU) cores on a chip and thenumber of CPU cores connected to a shared memory continues to growsignificantly to support growing workload capacity demand. Theincreasing number of CPUs cooperating to process the same workloads putsa significant burden on software scalability, for example, shared queuesor data structures protected by traditional semaphores become hot spotsand lead to sub-linear n-way scaling curves. Traditionally this has beencountered by implementing finer-grained locking in software, and withlower latency/higher bandwidth interconnects in hardware. Implementingfine-grained locking to improve software scalability can be verycomplicated and error-prone, and at today's CPU frequencies, thelatencies of hardware interconnects are limited by the physicaldimension of the chips and systems, and by the speed of light.

Implementations of hardware Transactional Memory (HTM, or in thisdiscussion, simply TM) have been introduced, wherein a group ofinstructions—called a transaction—operate in an atomic manner on a datastructure in memory, as viewed by other central processing units (CPUs)and the I/O subsystem (atomic operation is also known as “blockconcurrent” or “serialized” in other literature). The transactionexecutes optimistically without obtaining a lock, but may need to abortand retry the transaction execution if an operation, of the executingtransaction, on a memory location conflicts with another operation onthe same memory location. Previously, software transactional memoryimplementations have been proposed to support software TransactionalMemory (TM). However, hardware TM can provide improved performanceaspects and ease of use over software TM.

U.S. Patent Application Publication US20080244354 A1 titled “Apparatusand method for redundant multi-threading with recovery” filed 2007 Mar.28 and incorporated by reference herein teaches a method and apparatusfor reducing the effect of soft errors in a computer system is provided.Soft errors are detected by combining software redundant threading andinstruction duplication. Upon detection of a soft error, errors arerecovered through the use of software check pointing/rollbacktechnology. Reliable regions are identified by vulnerability profilingand redundant multi-threading is applied to the identified reliableregions.

U.S. Patent Application Publication US20120210162 A1 titled “Staterecovery and lockstep execution restart in a system with multiprocessorpairing” filed 2011 Feb. 15 and incorporated by reference herein teachesa system, method and computer program product for a multiprocessingsystem to offer selective pairing of processor cores for increasedprocessing reliability. A selective pairing facility is provided thatselectively connects, i.e., pairs, multiple microprocessor or processorcores to provide one highly reliable thread (or thread group). Eachpaired microprocessor or processor cores that provide one highlyreliable thread for high-reliability connect with a system componentssuch as a memory “nest” (or memory hierarchy), an optional systemcontroller, and optional interrupt controller, optional I/O orperipheral devices, etc. The memory nest is attached to a selectivepairing facility via a switch or a bus. Each selectively pairedprocessor core is includes a transactional execution facility, whereinthe system is configured to enable processor rollback to a previousstate and reinitialize lockstep execution in order to recover from anincorrect execution when an incorrect execution has been detected by theselective pairing facility.

SUMMARY

Aspects of an embodiment of the present invention disclose a method,computer program product, and computing system for resolving terminatedtransactions in a transactional memory environment. A processorinitiates a hardware transaction in a computing environment, wherein thehardware transaction accesses a memory location, and wherein thehardware transaction includes a transaction begin indicator and atransaction end indicator. A processor detects a conflicting access ofthe memory location while executing the hardware transaction. Aprocessor aborts the hardware transaction based on the conflictingaccess of the memory location. Hardware determines that the conflictingaccess of the memory location is a transient condition. A processorreinitiates the hardware transaction.

BRIEF DESCRIPTION OF THE DRAWINGS

FIG. 1 depicts an example multicore transactional memory environment, inaccordance with an illustrative embodiment.

FIG. 2 depicts an example multicore transactional memory environment, inaccordance with an illustrative embodiment.

FIG. 3 depicts example components of an example CPU, in accordance withan illustrative embodiment.

FIG. 4 is a block diagram of internal and external components of acomputing device, in accordance with one embodiment of the presentinvention.

FIG. 5 is a flowchart of the steps of an approach for processing ahardware transaction and identifying that an interference has occurred,in accordance with one embodiment of the present invention.

FIG. 6 is a block diagram of regulation logic 420 and associatedmodules, in accordance with one embodiment of the present invention.

FIG. 7 depicts a flowchart of the steps of regulation logic fordetecting whether a transient condition was the cause of an abortedhardware transaction and determining whether to retry the hardwaretransaction, in accordance with one embodiment of the present invention.

FIG. 8 depicts a flowchart of the steps of regulation logic forselecting retry parameters, in accordance with one embodiment of thepresent invention.

DETAILED DESCRIPTION

A transaction within a computer program or computer applicationcomprises program instructions performing multiple store operations thatappear to run and complete as a single, atomic operation. The programinstructions forming a current transaction comprise a transaction beginindicator, a plurality of instructions (e.g., arithmetic, load, branchor store operations), and a transaction end indicator. A near-end oftransaction indicator is triggered based on a speculative look aheadoperation, and enabling near-end-transaction processing mode, such thatan interfering operation may be delayed to allow the current transactionto complete. A halt operation, also referred to as an abort operation,as used herein refers to an operation responsive to a condition wheretwo transactions have been detected to interfere where at least onetransaction must be aborted and the state of the processor is reset tothe state at the beginning of the aborted transaction by performing arollback. This Detailed Description section is divided into thefollowing subsections: (i) The Hardware and Software Environment; (ii)Example Embodiment; (iii) Further Comments and/or Embodiments; and (iv)Definitions.

I. The Hardware and Software Environments

The present invention may be a system, a method, and/or a computerprogram product. The computer program product may include a computerreadable storage medium (or media) having computer readable programinstructions thereon for causing a processor to carry out aspects of thepresent invention.

A. Transaction Execution Environment

Historically, a computer system or processor had only a single processor(aka processing unit or central processing unit). The processor includedan instruction processing unit (IPU), a branch unit, a memory controlunit and the like. Such processors were capable of executing a singlethread of a program at a time. Operating systems were developed thatcould time-share a processor by dispatching a program to be executed onthe processor for a period of time, and then dispatching another programto be executed on the processor for another period of time. Astechnology evolved, memory subsystem caches were often added to theprocessor as well as complex dynamic address translation includingtranslation lookaside buffers (TLBs). The IPU itself was often referredto as a processor. As technology continued to evolve, an entireprocessor could be packaged on a single semiconductor chip or die, sucha processor was referred to as a microprocessor. Then processors weredeveloped that incorporated multiple IPUs, such processors were oftenreferred to as multi-processors. Each such processor of amulti-processor computer system (processor) may include individual orshared caches, memory interfaces, system bus, address translationmechanism and the like. Virtual machine and instruction set architecture(ISA) emulators added a layer of software to a processor, that providedthe virtual machine with multiple “virtual processors” (aka processors)by time-slice usage of a single IPU in a single hardware processor. Astechnology further evolved, multi-threaded processors were developed,enabling a single hardware processor having a single multi-thread IPU toprovide a capability of simultaneously executing threads of differentprograms, thus each thread of a multi-threaded processor appeared to theoperating system as a processor. As technology further evolved, it waspossible to put multiple processors (each having an IPU) on a singlesemiconductor chip or die. These processors were referred to processorcores or just cores. Thus the terms such as processor, centralprocessing unit, processing unit, microprocessor, core, processor core,processor thread, and thread, for example, are often usedinterchangeably. Aspects of embodiments herein may be practiced by anyor all processors including those shown supra, without departing fromthe teachings herein. Wherein the term “thread” or “processor thread” isused herein, it is expected that particular advantage of the embodimentmay be had in a processor thread implementation.

Transaction Execution in Intel Based Embodiments

In “Intel Architecture Instruction Set Extensions Programming Reference”319433-012A, February 2012, incorporated herein by reference in itsentirety, Chapter 8 teaches, in part, that multithreaded applicationsmay take advantage of increasing numbers of CPU cores to achieve higherperformance. (Note: the term “Intel” may be subject to trademark rightsin various jurisdictions throughout the world and are used here only inreference to the products or services properly denominated by the marksto the extent that such trademark rights may exist.) However, thewriting of multi-threaded applications requires programmers tounderstand and take into account data sharing among the multiplethreads. Access to shared data typically requires synchronizationmechanisms. These synchronization mechanisms are used to ensure thatmultiple threads update shared data by serializing operations that areapplied to the shared data, often through the use of a critical sectionthat is protected by a lock. Since serialization limits concurrency,programmers try to limit the overhead due to synchronization.

Intel Transactional Synchronization Extensions (Intel TSX) allow aprocessor to dynamically determine whether threads need to be serializedthrough lock-protected critical sections, and to perform thatserialization only when required. (Note: the term(s) “Intel,” “TSX,”and/or “Intel TSX” may be subject to trademark rights in variousjurisdictions throughout the world and are used here only in referenceto the products or services properly denominated by the marks to theextent that such trademark rights may exist.) This allows the processorto expose and exploit concurrency that is hidden in an applicationbecause of dynamically unnecessary synchronization.

With Intel TSX, programmer-specified code regions (also referred to as“transactional regions” or just “transactions”) are executedtransactionally. If the transactional execution completes successfully,then all memory operations performed within the transactional regionwill appear to have occurred instantaneously when viewed from otherprocessors. A processor makes the memory operations of the executedtransaction, performed within the transactional region, visible to otherprocessors only when a successful commit occurs, i.e., when thetransaction successfully completes execution. This process is oftenreferred to as an atomic commit.

Intel TSX provides two software interfaces to specify regions of codefor transactional execution. Hardware Lock Elision (HLE) is a legacycompatible instruction set extension (comprising the XACQUIRE andXRELEASE prefixes) to specify transactional regions. RestrictedTransactional Memory (RTM) is a new instruction set interface(comprising the XBEGIN, XEND, and XABORT instructions) for programmersto define transactional regions in a more flexible manner than thatpossible with HLE. HLE is for programmers who prefer the backwardcompatibility of the conventional mutual exclusion programming model andwould like to run HLE-enabled software on legacy hardware but would alsolike to take advantage of the new lock elision capabilities on hardwarewith HLE support. RTM is for programmers who prefer a flexible interfaceto the transactional execution hardware. In addition, Intel TSX alsoprovides an XTEST instruction. This instruction allows software to querywhether the logical processor is transactionally executing in atransactional region identified by either HLE or RTM.

Since a successful transactional execution ensures an atomic commit, theprocessor executes the code region optimistically without explicitsynchronization. If synchronization was unnecessary for that specificexecution, execution can commit without any cross-thread serialization.If the processor cannot commit atomically, then the optimistic executionfails. When this happens, the processor will roll back the execution, aprocess referred to as a transactional abort. On a transactional abort,the processor will discard all updates performed in the memory regionused by the transaction, restore architectural state to appear as if theoptimistic execution never occurred, and resume executionnon-transactionally.

A processor can perform a transactional abort for numerous reasons. Aprimary reason to abort a transaction is due to conflicting memoryaccesses between the transactionally executing logical processor andanother logical processor. Such conflicting memory accesses may preventa successful transactional execution. Memory addresses read from withina transactional region constitute the read-set of the transactionalregion and addresses written to within the transactional regionconstitute the write-set of the transactional region. Intel TSXmaintains the read- and write-sets at the granularity of a cache line. Aconflicting memory access occurs if another logical processor eitherreads a location that is part of the transactional region's write-set orwrites a location that is a part of either the read- or write-set of thetransactional region. A conflicting access typically means thatserialization is required for this code region. Since Intel TSX detectsdata conflicts at the granularity of a cache line, unrelated datalocations placed in the same cache line will be detected as conflictsthat result in transactional aborts. Transactional aborts may also occurdue to limited transactional resources. For example, the amount of dataaccessed in the region may exceed an implementation-specific capacity.Additionally, some instructions and system events may causetransactional aborts. Frequent transactional aborts result in wastedcycles and increased inefficiency.

Hardware Lock Elision

Hardware Lock Elision (HLE) provides a legacy compatible instruction setinterface for programmers to use transactional execution. HLE providestwo new instruction prefix hints: XACQUIRE and XRELEASE.

With HLE, a programmer adds the XACQUIRE prefix to the front of theinstruction that is used to acquire the lock that is protecting thecritical section. The processor treats the prefix as a hint to elide thewrite associated with the lock acquire operation. Even though the lockacquire has an associated write operation to the lock, the processordoes not add the address of the lock to the transactional region'swrite-set nor does it issue any write requests to the lock. Instead, theaddress of the lock is added to the read-set. The logical processorenters transactional execution. If the lock was available before theXACQUIRE prefixed instruction, then all other processors will continueto see the lock as available afterwards. Since the transactionallyexecuting logical processor neither added the address of the lock to itswrite-set nor performed externally visible write operations to the lock,other logical processors can read the lock without causing a dataconflict. This allows other logical processors to also enter andconcurrently execute the critical section protected by the lock. Theprocessor automatically detects any data conflicts that occur during thetransactional execution and will perform a transactional abort ifnecessary.

Even though the eliding processor did not perform any external writeoperations to the lock, the hardware ensures program order of operationson the lock. If the eliding processor itself reads the value of the lockin the critical section, it will appear as if the processor had acquiredthe lock, i.e. the read will return the non-elided value. This behaviorallows an HLE execution to be functionally equivalent to an executionwithout the HLE prefixes.

An XRELEASE prefix can be added in front of an instruction that is usedto release the lock protecting a critical section. Releasing the lockinvolves a write to the lock. If the instruction is to restore the valueof the lock to the value the lock had prior to the XACQUIRE prefixedlock acquire operation on the same lock, then the processor elides theexternal write request associated with the release of the lock and doesnot add the address of the lock to the write-set. The processor thenattempts to commit the transactional execution.

With HLE, if multiple threads execute critical sections protected by thesame lock but they do not perform any conflicting operations on eachother's data, then the threads can execute concurrently and withoutserialization. Even though the software uses lock acquisition operationson a common lock, the hardware recognizes this, elides the lock, andexecutes the critical sections on the two threads without requiring anycommunication through the lock—if such communication was dynamicallyunnecessary.

If the processor is unable to execute the region transactionally, thenthe processor will execute the region non-transactionally and withoutelision. HLE enabled software has the same forward progress guaranteesas the underlying non-HLE lock-based execution. For successful HLEexecution, the lock and the critical section code must follow certainguidelines. These guidelines only affect performance; and failure tofollow these guidelines will not result in a functional failure.Hardware without HLE support will ignore the XACQUIRE and XRELEASEprefix hints and will not perform any elision since these prefixescorrespond to the REPNE/REPE IA-32 prefixes which are ignored on theinstructions where XACQUIRE and XRELEASE are valid. Importantly, HLE iscompatible with the existing lock-based programming model. Improper useof hints will not cause functional bugs though it may expose latent bugsalready in the code.

Restricted Transactional Memory (RTM) provides a flexible softwareinterface for transactional execution. RTM provides three newinstructions—XBEGIN, XEND, and XABORT—for programmers to start, commit,and abort a transactional execution.

The programmer uses the XBEGIN instruction to specify the start of atransactional code region and the XEND instruction to specify the end ofthe transactional code region. If the RTM region could not besuccessfully executed transactionally, then the XBEGIN instruction takesan operand that provides a relative offset to the fallback instructionaddress.

A processor may abort RTM transactional execution for many reasons. Inmany instances, the hardware automatically detects transactional abortconditions and restarts execution from the fallback instruction addresswith the architectural state corresponding to that present at the startof the XBEGIN instruction and the EAX register updated to describe theabort status.

The XABORT instruction allows programmers to abort the execution of anRTM region explicitly. The XABORT instruction takes an 8-bit immediateargument that is loaded into the EAX register and will thus be availableto software following an RTM abort. RTM instructions do not have anydata memory location associated with them. While the hardware providesno guarantees as to whether an RTM region will ever successfully committransactionally, most transactions that follow the recommendedguidelines are expected to successfully commit transactionally. However,programmers must always provide an alternative code sequence in thefallback path to guarantee forward progress. This may be as simple asacquiring a lock and executing the specified code regionnon-transactionally. Further, a transaction that always aborts on agiven implementation may complete transactionally on a futureimplementation. Therefore, programmers must ensure the code paths forthe transactional region and the alternative code sequence arefunctionally tested.

Detection of HLE Support

A processor supports HLE execution if CPUID.07H.EBX.HLE [bit 4]=1.However, an application can use the HLE prefixes (XACQUIRE and XRELEASE)without checking whether the processor supports HLE. Processors withoutHLE support ignore these prefixes and will execute the code withoutentering transactional execution.

Detection of RTM Support

A processor supports RTM execution if CPUID.07H.EBX.RTM [bit 11]=1. Anapplication must check if the processor supports RTM before it uses theRTM instructions (XBEGIN, XEND, XABORT). These instructions willgenerate a #UD exception when used on a processor that does not supportRTM.

Detection of XTEST Instruction

A processor supports the XTEST instruction if it supports either HLE orRTM. An application must check either of these feature flags beforeusing the XTEST instruction. This instruction will generate a #UDexception when used on a processor that does not support either HLE orRTM.

Querying Transactional Execution Status

The XTEST instruction can be used to determine the transactional statusof a transactional region specified by HLE or RTM. Note, while the HLEprefixes are ignored on processors that do not support HLE, the XTESTinstruction will generate a #UD exception when used on processors thatdo not support either HLE or RTM.

Requirements for HLE Locks

For HLE execution to successfully commit transactionally, the lock mustsatisfy certain properties and access to the lock must follow certainguidelines.

An XRELEASE prefixed instruction must restore the value of the elidedlock to the value it had before the lock acquisition. This allowshardware to safely elide locks by not adding them to the write-set. Thedata size and data address of the lock release (XRELEASE prefixed)instruction must match that of the lock acquire (XACQUIRE prefixed) andthe lock must not cross a cache line boundary.

Software should not write to the elided lock inside a transactional HLEregion with any instruction other than an XRELEASE prefixed instruction,otherwise such a write may cause a transactional abort. In addition,recursive locks (where a thread acquires the same lock multiple timeswithout first releasing the lock) may also cause a transactional abort.Note that software can observe the result of the elided lock acquireinside the critical section. Such a read operation will return the valueof the write to the lock.

The processor automatically detects violations to these guidelines, andsafely transitions to a non-transactional execution without elision.Since Intel TSX detects conflicts at the granularity of a cache line,writes to data collocated on the same cache line as the elided lock maybe detected as data conflicts by other logical processors eliding thesame lock.

Transactional Nesting

Both HLE and RTM support nested transactional regions. However, atransactional abort restores state to the operation that startedtransactional execution: either the outermost XACQUIRE prefixed HLEeligible instruction or the outermost XBEGIN instruction. The processortreats all nested transactions as one transaction.

HLE Nesting and Elision

Programmers can nest HLE regions up to an implementation specific depthof MAX_HLE_NEST_COUNT. Each logical processor tracks the nesting countinternally but this count is not available to software. An XACQUIREprefixed HLE-eligible instruction increments the nesting count, and anXRELEASE prefixed HLE-eligible instruction decrements it. The logicalprocessor enters transactional execution when the nesting count goesfrom zero to one. The logical processor attempts to commit only when thenesting count becomes zero. A transactional abort may occur if thenesting count exceeds MAX_HLE_NEST_COUNT.

In addition to supporting nested HLE regions, the processor can alsoelide multiple nested locks. The processor tracks a lock for elisionbeginning with the XACQUIRE prefixed HLE eligible instruction for thatlock and ending with the XRELEASE prefixed HLE eligible instruction forthat same lock. The processor can, at any one time, track up to aMAX_HLE_ELIDED_LOCKS number of locks. For example, if the implementationsupports a MAX_HLE_ELIDED_LOCKS value of two and if the programmer neststhree HLE identified critical sections (by performing XACQUIRE prefixedHLE eligible instructions on three distinct locks without performing anintervening XRELEASE prefixed HLE eligible instruction on any one of thelocks), then the first two locks will be elided, but the third won't beelided (but will be added to the transaction's write-set). However, theexecution will still continue transactionally. Once an XRELEASE for oneof the two elided locks is encountered, a subsequent lock acquiredthrough the XACQUIRE prefixed HLE eligible instruction will be elided.

The processor attempts to commit the HLE execution when all elidedXACQUIRE and XRELEASE pairs have been matched, the nesting count goes tozero, and the locks have satisfied requirements. If execution cannotcommit atomically, then execution transitions to a non-transactionalexecution without elision as if the first instruction did not have anXACQUIRE prefix.

RTM Nesting

Programmers can nest RTM regions up to an implementation specificMAX_RTM_NEST_COUNT. The logical processor tracks the nesting countinternally but this count is not available to software. An XBEGINinstruction increments the nesting count, and an XEND instructiondecrements the nesting count. The logical processor attempts to commitonly if the nesting count becomes zero. A transactional abort occurs ifthe nesting count exceeds MAX_RTM_NEST_COUNT.

Nesting HLE and RTM

HLE and RTM provide two alternative software interfaces to a commontransactional execution capability. Transactional processing behavior isimplementation specific when HLE and RTM are nested together, e.g., HLEis inside RTM or RTM is inside HLE. However, in all cases, theimplementation will maintain HLE and RTM semantics. An implementationmay choose to ignore HLE hints when used inside RTM regions, and maycause a transactional abort when RTM instructions are used inside HLEregions. In the latter case, the transition from transactional tonon-transactional execution occurs seamlessly since the processor willre-execute the HLE region without actually doing elision, and thenexecute the RTM instructions.

Abort Status Definition

RTM uses the EAX register to communicate abort status to software.Following an RTM abort the EAX register has the following definition.

TABLE 1 RTM Abort Status Definition EAX Register Bit Position Meaning 0Set if abort caused by XABORT instruction 1 If set, the transaction maysucceed on retry, this bit is always clear if bit 0 is set 2 Set ifanother logical processor conflicted with a memory address that was partof the transaction that aborted 3 Set if an internal buffer overflowed 4Set if a debug breakpoint was hit 5 Set if an abort occurred duringexecution of a nested transaction 23:6 Reserved 31-24 XABORT argument(only valid if bit 0 set, otherwise reserved)

The EAX abort status for RTM only provides causes for aborts. It doesnot by itself encode whether an abort or commit occurred for the RTMregion. The value of EAX can be 0 following an RTM abort. For example, aCPUID instruction when used inside an RTM region causes a transactionalabort and may not satisfy the requirements for setting any of the EAXbits. This may result in an EAX value of 0.

RTM Memory Ordering

A successful RTM commit causes all memory operations in the RTM regionto appear to execute atomically. A successfully committed RTM regionconsisting of an XBEGIN followed by an XEND, even with no memoryoperations in the RTM region, has the same ordering semantics as a LOCKprefixed instruction.

The XBEGIN instruction does not have fencing semantics. However, if anRTM execution aborts, then all memory updates from within the RTM regionare discarded and are not made visible to any other logical processor.

RTM-Enabled Debugger Support

By default, any debug exception inside an RTM region will cause atransactional abort and will redirect control flow to the fallbackinstruction address with architectural state recovered and bit 4 in EAXset. However, to allow software debuggers to intercept execution ondebug exceptions, the RTM architecture provides additional capability.

If bit 11 of DR7 and bit 15 of the IA32_DEBUGCTL_MSR are both 1, any RTMabort due to a debug exception (#DB) or breakpoint exception (#BP)causes execution to roll back and restart from the XBEGIN instructioninstead of the fallback address. In this scenario, the EAX register willalso be restored back to the point of the XBEGIN instruction.

Programming Considerations

Typical programmer-identified regions are expected to transactionallyexecute and commit successfully. However, Intel TSX does not provide anysuch guarantee. A transactional execution may abort for many reasons. Totake full advantage of the transactional capabilities, programmersshould follow certain guidelines to increase the probability of theirtransactional execution committing successfully.

This section discusses various events that may cause transactionalaborts. The architecture ensures that updates performed within atransaction that subsequently aborts execution will never becomevisible. Only committed transactional executions initiate an update tothe architectural state. Transactional aborts never cause functionalfailures and only affect performance.

Instruction Based Considerations

Programmers can use any instruction safely inside a transaction (HLE orRTM) and can use transactions at any privilege level. However, someinstructions will always abort the transactional execution and causeexecution to seamlessly and safely transition to a non-transactionalpath.

Intel TSX allows for most common instructions to be used insidetransactions without causing aborts. The following operations inside atransaction do not typically cause an abort:

-   -   Operations on the instruction pointer register, general purpose        registers (GPRs) and the status flags (CF, OF, SF, PF, AF, and        ZF); and    -   Operations on XMM and YMM registers and the MXCSR register.

However, programmers must be careful when intermixing SSE and AVXoperations inside a transactional region. Intermixing SSE instructionsaccessing XMM registers and AVX instructions accessing YMM registers maycause transactions to abort. Programmers may use REP/REPNE prefixedstring operations inside transactions. However, long strings may causeaborts. Further, the use of CLD and STD instructions may cause aborts ifthey change the value of the DF flag. However, if DF is 1, the STDinstruction will not cause an abort. Similarly, if DF is 0, then the CLDinstruction will not cause an abort.

Instructions not enumerated here as causing abort when used inside atransaction will typically not cause a transaction to abort (examplesinclude but are not limited to MFENCE, LFENCE, SFENCE, RDTSC, RDTSCP,etc.).

The following instructions will abort transactional execution on anyimplementation:

-   -   XABORT    -   CPUID    -   PAUSE

In addition, in some implementations, the following instructions mayalways cause transactional aborts. These instructions are not expectedto be commonly used inside typical transactional regions. However,programmers must not rely on these instructions to force a transactionalabort, since whether they cause transactional aborts is implementationdependent.

-   -   Operations on X87 and MMX architecture state. This includes all        MMX and X87 instructions, including the FXRSTOR and FXSAVE        instructions.    -   Update to non-status portion of EFLAGS: CLI, STI, POPFD, POPFQ,        CLTS.    -   Instructions that update segment registers, debug registers        and/or control registers:    -   MOV to DS/ES/FS/GS/SS, POP DS/ES/FS/GS/SS, LDS, LES, LFS, LGS,        LSS, SWAPGS, WRFSBASE, WRGSBASE, LGDT, SGDT, LIDT, SIDT, LLDT,        SLDT, LTR, STR, Far CALL, Far JMP, Far RET, IRET, MOV to DRx,        MOV to CR0/CR2/CR3/CR4/CR8 and LMSW.    -   Ring transitions: SYSENTER, SYSCALL, SYSEXIT, and SYSRET.    -   TLB and Cacheability control: CLFLUSH, INVD, WBINVD, INVLPG,        INVPCID, and memory instructions with a non-temporal hint        (MOVNTDQA, MOVNTDQ, MOVNTI, MOVNTPD, MOVNTPS, and MOVNTQ).    -   Processor state save: XSAVE, XSAVEOPT, and XRSTOR.    -   Interrupts: INTn, INTO.    -   IO: IN, INS, REP INS, OUT, OUTS, REP OUTS and their variants.    -   VMX: VMPTRLD, VMPTRST, VMCLEAR, VMREAD, VMWRITE, VMCALL,        VMLAUNCH, VMRESUME, VMXOFF, VMXON, INVEPT, and INVVPID.    -   SMX: GETSEC.    -   UD2, RSM, RDMSR, WRMSR, HLT, MONITOR, MWAIT, XSETBV, VZEROUPPER,        MASKMOVQ, and V/MASKMOVDQU.        Runtime Considerations

In addition to the instruction-based considerations, runtime events maycause transactional execution to abort. These may be due to data accesspatterns or micro-architectural implementation features. The followinglist is not a comprehensive discussion of all abort causes.

Any fault or trap in a transaction that must be exposed to software willbe suppressed. Transactional execution will abort and execution willtransition to a non-transactional execution, as if the fault or trap hadnever occurred. If an exception is not masked, then that un-maskedexception will result in a transactional abort and the state will appearas if the exception had never occurred.

Synchronous exception events (#DE, #OF, #NP, #SS, #GP, #BR, #UD, #AC,#XF, #PF, #NM, #TS, #MF, #DB, #BP/INT3) that occur during transactionalexecution may cause an execution not to commit transactionally, andrequire a non-transactional execution. These events are suppressed as ifthey had never occurred. With HLE, since the non-transactional code pathis identical to the transactional code path, these events will typicallyreappear when the instruction that caused the exception is re-executednon-transactionally, causing the associated synchronous events to bedelivered appropriately in the non-transactional execution. Asynchronousevents (NMI, SMI, INTR, IPI, PMI, etc.) occurring during transactionalexecution may cause the transactional execution to abort and transitionto a non-transactional execution. The asynchronous events will be pendedand handled after the transactional abort is processed.

Transactions only support write-back cacheable memory type operations. Atransaction may always abort if the transaction includes operations onany other memory type. This includes instruction fetches to UC memorytype.

Memory accesses within a transactional region may require the processorto set the Accessed and Dirty flags of the referenced page table entry.The behavior of how the processor handles this is implementationspecific. Some implementations may allow the updates to these flags tobecome externally visible even if the transactional region subsequentlyaborts. Some Intel TSX implementations may choose to abort thetransactional execution if these flags need to be updated. Further, aprocessor's page-table walk may generate accesses to its owntransactionally written but uncommitted state. Some Intel TSXimplementations may choose to abort the execution of a transactionalregion in such situations. Regardless, the architecture ensures that, ifthe transactional region aborts, then the transactionally written statewill not be made architecturally visible through the behavior ofstructures such as TLBs.

Executing self-modifying code transactionally may also causetransactional aborts. Programmers must continue to follow the Intelrecommended guidelines for writing self-modifying and cross-modifyingcode even when employing HLE and RTM. While an implementation of RTM andHLE will typically provide sufficient resources for executing commontransactional regions, implementation constraints and excessive sizesfor transactional regions may cause a transactional execution to abortand transition to a non-transactional execution. The architectureprovides no guarantee of the amount of resources available to dotransactional execution and does not guarantee that a transactionalexecution will ever succeed.

Conflicting requests to a cache line accessed within a transactionalregion may prevent the transaction from executing successfully. Forexample, if logical processor P0 reads line A in a transactional regionand another logical processor P1 writes line A (either inside or outsidea transactional region) then logical processor P0 may abort if logicalprocessor P1's write interferes with processor P0's ability to executetransactionally.

Similarly, if P0 writes line A in a transactional region and P1 reads orwrites line A (either inside or outside a transactional region), then P0may abort if P1's access to line A interferes with P0's ability toexecute transactionally. In addition, other coherence traffic may attimes appear as conflicting requests and may cause aborts. While thesefalse conflicts may happen, they are expected to be uncommon. Theconflict resolution policy to determine whether P0 or P1 aborts in theabove scenarios is implementation specific.

Generic Transaction Execution Embodiments:

According to “ARCHITECTURES FOR TRANSACTIONAL MEMORY”, a dissertationsubmitted to the Department of Computer Science and the Committee onGraduate Studies of Stanford University in partial fulfillment of therequirements for the Degree of Doctor of Philosophy, by Austen McDonald,June 2009, incorporated by reference herein in its entirety,fundamentally, there are three mechanisms needed to implement an atomicand isolated transactional region: versioning, conflict detection, andcontention management.

To make a transactional code region appear atomic, all the modificationsperformed by that transactional code region must be stored and keptisolated from other transactions until commit time. The system does thisby implementing a versioning policy. Two versioning paradigms exist:eager and lazy. An eager versioning system stores newly generatedtransactional values in place and stores previous memory values on theside, in what is called an undo-log. A lazy versioning system stores newvalues temporarily in what is called a write buffer, copying them tomemory only on commit. In either system, the cache is used to optimizestorage of new versions.

To ensure that transactions appear to be performed atomically, conflictsmust be detected and resolved. The two systems, i.e., the eager and lazyversioning systems, detect conflicts by implementing a conflictdetection policy, either optimistic or pessimistic. An optimistic systemexecutes transactions in parallel, checking for conflicts only when atransaction commits. A pessimistic system checks for conflicts at eachload and store. Similar to versioning, conflict detection also uses thecache, marking each line as either part of the read-set, part of thewrite-set, or both. The two systems resolve conflicts by implementing acontention management policy. Many contention management policies exist,some are more appropriate for optimistic conflict detection and some aremore appropriate for pessimistic. Described below are some examplepolicies.

Since each transactional memory (TM) system needs both versioningdetection and conflict detection, these options give rise to fourdistinct TM designs: Eager-Pessimistic (EP), Eager-Optimistic (EO),Lazy-Pessimistic (LP), and Lazy-Optimistic (LO). Table 2 brieflydescribes all four distinct TM designs.

FIGS. 1 and 2 depict an example of a multicore TM environment. FIG. 1shows many TM-enabled CPUs (CPU1 114 a, CPU2 114 b, and other CPUs notshown) on one die 100, connected with an interconnect 122, undermanagement of an interconnect control 120 a, 120 b. Each CPU 114 a and114 b (also known as a Processor) may have a split cache consisting ofInstruction Cache 116 a and 116 b for caching instructions from memoryto be executed and Data Cache with TM support 118 a and 118 b forcaching data (operands) of memory locations to be operated on by CPU 114a and 114 b (in FIG. 1, each CPU 114 a, 114 b and its associated cachesare referenced as 112 a and 112 b). In an implementation, caches ofmultiple dies 100 are interconnected to support cache coherency betweenthe caches of multiple dies 100. In an implementation, a single cache,rather than the split cache is employed holding both instructions anddata. In implementations, the CPU caches are one level of caching in ahierarchical cache structure. For example each die 100 may employ sharedcache 124 to be shared amongst all the CPUs on die 100. In anotherimplementation, each die may have access to shared cache 124, sharedamongst all the processors of all dies 100.

FIG. 2 shows the details of transactional CPU environment 112 a, havingCPU 114 a, including additions to support TM. Transactional CPU(processor) 114 a may include hardware for supporting RegisterCheckpoints 126 and special TM Registers 128. The transactional CPUcache may have MESI bits 130, Tags 140 and Data 142 of a conventionalcache but also, for example, R bits 132 showing a line has been read byCPU 114 a while executing a transaction and W bits 138 showing a linehas been written-to by CPU 114 a while executing a transaction.

A key detail for programmers in any TM system is how non-transactionalaccesses interact with transactions. By design, transactional accessesare screened from each other using the mechanisms above. However, theinteraction between a regular, non-transactional load with a transactioncontaining a new value for that address must still be considered. Inaddition, the interaction between a non-transactional store with atransaction that has read that address must also be explored. These areissues of the database concept isolation.

A TM system is said to implement strong isolation, sometimes calledstrong atomicity, when every non-transactional load and store acts likean atomic transaction. Therefore, non-transactional loads cannot seeuncommitted data and non-transactional stores cause atomicity violationsin any transactions that have read that address. A system where this isnot the case is said to implement weak isolation, sometimes called weakatomicity.

Strong isolation is often more desirable than weak isolation due to therelative ease of conceptualization and implementation of strongisolation. Additionally, if a programmer has forgotten to surround someshared memory references with transactions, causing bugs, then withstrong isolation, the programmer will often detect that oversight usinga simple debug interface because the programmer will see anon-transactional region causing atomicity violations. Also, programswritten in one model may work differently on another model. Further,strong isolation is often easier to support in hardware TM than weakisolation. With strong isolation, since the coherence protocol alreadymanages load and store communication between processors, transactionscan detect non-transactional loads and stores and act appropriately. Toimplement strong isolation in software Transactional Memory (TM),non-transactional code must be modified to include read- andwrite-barriers; potentially crippling performance. Although great efforthas been expended to remove many unneeded barriers, such techniques areoften complex and performance is typically far lower than that of HTMs.

TABLE 2 Transactional Memory Design Space VERSIONING Lazy Eager CONFLICTOptimistic Storing updates in a write Not practical: waiting to updateDETECTION buffer; detecting conflicts at memory until commit time butcommit time. detecting conflicts at access time guarantees wasted workand provides no advantage. Pessimistic Storing updates in a writeUpdating memory, keeping old buffer; detecting conflicts at values inundo log; detecting access time. conflicts at access time.

Table 2 illustrates the fundamental design space of transactional memory(versioning and conflict detection).

Eager-Pessimistic (EP)

This first TM design described below is known as Eager-Pessimistic. AnEP system stores its write-set “in place” (hence the name “eager”) and,to support rollback, stores the old values of overwritten lines in an“undo log”. Processors use the W 138 and R 132 cache bits to track readand write-sets and detect conflicts when receiving snooped loadrequests. Perhaps the most notable examples of EP systems in knownliterature are LogTM and UTM.

Beginning a transaction in an EP system is much like beginning atransaction in other systems: tm_begin( ) takes a register checkpoint,and initializes any status registers. An EP system also requiresinitializing the undo log, the details of which are dependent on the logformat, but often involve initializing a log base pointer to a region ofpre-allocated, thread-private memory, and clearing a log boundsregister.

Versioning: In EP, due to the way eager versioning is designed tofunction, the MESI 130 state transitions (cache line indicatorscorresponding to Modified, Exclusive, Shared, and Invalid code states)are left mostly unchanged. Outside of a transaction, the MESI 130 statetransitions are left completely unchanged. When reading a line inside atransaction, the standard coherence transitions apply (S (Shared)→S, I(Invalid)→S, or I→E (Exclusive)), issuing a load miss as needed, but theR 132 bit is also set. Likewise, writing a line applies the standardtransitions (S→M, E→I, I→M), issuing a miss as needed, but also sets theW 138 (Written) bit. The first time a line is written, the old versionof the entire line is loaded then written to the undo log to preserve itin case the current transaction aborts. The newly written data is thenstored “in-place,” over the old data.

Conflict Detection: Pessimistic conflict detection uses coherencemessages exchanged on misses, or upgrades, to look for conflicts betweentransactions. When a read miss occurs within a transaction, otherprocessors receive a load request; but they ignore the request if theydo not have the needed line. If the other processors have the neededline non-speculatively or have the line R 132 (Read), they downgradethat line to S, and in certain cases issue a cache-to-cache transfer ifthey have the line in MESI's 130 M or E state. However, if the cache hasthe line W 138, then a conflict is detected between the two transactionsand additional action(s) must be taken.

Similarly, when a transaction seeks to upgrade a line from shared tomodified (on a first write), the transaction issues an exclusive loadrequest, which is also used to detect conflicts. If a receiving cachehas the line non-speculatively, then the line is invalidated, and incertain cases a cache-to-cache transfer (M or E states) is issued. But,if the line is R 132 or W 138, a conflict is detected.

Validation: Because conflict detection is performed on every load, atransaction always has exclusive access to its own write-set. Therefore,validation does not require any additional work.

Commit: Since eager versioning stores the new version of data items inplace, the commit process simply clears the W 138 and R 132 bits anddiscards the undo log.

Abort: When a transaction rolls back, the original version of each cacheline in the undo log must be restored, a process called “unrolling” or“applying” the log. This is done during tm_discard( ) and must be atomicwith regard to other transactions. Specifically, the write-set muststill be used to detect conflicts: this transaction has the only correctversion of lines in its undo log, and requesting transactions must waitfor the correct version to be restored from that log. Such a log can beapplied using a hardware state machine or software abort handler.

Eager-Pessimistic has the characteristics of: Commit is simple and sinceit is in-place, very fast. Similarly, validation is a no-op. Pessimisticconflict detection detects conflicts early, thereby reducing the numberof “doomed” transactions. For example, if two transactions are involvedin a Write-After-Read dependency, then that dependency is detectedimmediately in pessimistic conflict detection. However, in optimisticconflict detection such conflicts are not detected until the writercommits.

Eager-Pessimistic also has the characteristics of: As described above,the first time a cache line is written, the old value must be written tothe log, incurring extra cache accesses. Aborts are expensive as theyrequire undoing the log. For each cache line in the log, a load must beissued, perhaps going as far as main memory before continuing to thenext line. Pessimistic conflict detection also prevents certainserializable schedules from existing.

Additionally, because conflicts are handled as they occur, there is apotential for livelock and careful contention management mechanisms mustbe employed to guarantee forward progress.

Lazy-Optimistic (LO)

Another popular TM design is Lazy-Optimistic (LO), which stores itswrite-set in a “write buffer” or “redo log” and detects conflicts atcommit time (still using the R 132 and W 138 bits).

Versioning: Just as in the EP system, the MESI protocol of the LO designis enforced outside of the transactions. Once inside a transaction,reading a line incurs the standard MESI transitions but also sets the R132 bit. Likewise, writing a line sets the W 138 bit of the line, buthandling the MESI transitions of the LO design is different from that ofthe EP design. First, with lazy versioning, the new versions of writtendata are stored in the cache hierarchy until commit while othertransactions have access to old versions available in memory or othercaches. To make available the old versions, dirty lines (M lines) mustbe evicted when first written by a transaction. Second, no upgrademisses are needed because of the optimistic conflict detection feature:if a transaction has a line in the S state, it can simply write to itand upgrade that line to an M state without communicating the changeswith other transactions because conflict detection is done at committime.

Conflict Detection and Validation: To validate a transaction and detectconflicts, LO communicates the addresses of speculatively modified linesto other transactions only when it is preparing to commit. Onvalidation, the processor sends one, potentially large, network packetcontaining all the addresses in the write-set. Data is not sent, butleft in the cache of the committer and marked dirty (M). To build thispacket without searching the cache for lines marked W, a simple bitvector is used, called a “store buffer,” with one bit per cache line totrack these speculatively modified lines. Other transactions use thisaddress packet to detect conflicts: if an address is found in the cacheand the R 132 and/or W 138 bits are set, then a conflict is initiated.If the line is found but neither R 132 nor W 138 is set, then the lineis simply invalidated, which is similar to processing an exclusive load.

To support transaction atomicity, these address packets must be handledatomically, i.e., no two address packets may exist at once with the sameaddresses. In an LO system, this can be achieved by simply acquiring aglobal commit token before sending the address packet. However, atwo-phase commit scheme could be employed by first sending out theaddress packet, collecting responses, enforcing an ordering protocol(perhaps oldest transaction first), and committing once all responsesare satisfactory.

Commit: Once validation has occurred, commit needs no special treatment:simply clear W 138 and R 132 bits and the store buffer. Thetransaction's writes are already marked dirty in the cache and othercaches' copies of these lines have been invalidated via the addresspacket. Other processors can then access the committed data through theregular coherence protocol.

Abort: Rollback is equally easy: because the write-set is containedwithin the local caches, these lines can be invalidated, then clear W138 and R 132 bits and the store buffer. The store buffer allows W linesto be found to invalidate without the need to search the cache.

Lazy-Optimistic has the characteristics of: Aborts are very fast,requiring no additional loads or stores and making only local changes.More serializable schedules can exist than found in EP, which allows anLO system to more aggressively speculate that transactions areindependent, which can yield higher performance. Finally, the latedetection of conflicts can increase the likelihood of forward progress.

Lazy-Optimistic also has the characteristics of: Validation takes globalcommunication time proportional to size of write set. Doomedtransactions can waste work since conflicts are detected only at committime.

Lazy-Pessimistic (LP)

Lazy-Pessimistic (LP) represents a third TM design option, sittingsomewhere between EP and LO: storing newly written lines in a writebuffer but detecting conflicts on a per access basis.

Versioning: Versioning is similar but not identical to that of LO:reading a line sets its R bit 132, writing a line sets its W bit 138,and a store buffer is used to track W lines in the cache. Also, dirty(M) lines must be evicted when first written by a transaction, just asin LO. However, since conflict detection is pessimistic, load exclusivesmust be performed when upgrading a transactional line from I, S→M, whichis unlike LO.

Conflict Detection: LP's conflict detection operates the same as EP's:using coherence messages to look for conflicts between transactions.

Validation: Like in EP, pessimistic conflict detection ensures that atany point, a running transaction has no conflicts with any other runningtransaction, so validation is a no-op.

Commit: Commit needs no special treatment: simply clear W 138 and R 132bits and the store buffer, like in LO.

Abort: Rollback is also like that of LO: simply invalidate the write-setusing the store buffer and clear the W and R bits and the store buffer.

Eager-Optimistic (EO)

The LP has the characteristics of: Like LO, aborts are very fast. LikeEP, the use of pessimistic conflict detection reduces the number of“doomed” transactions. Like EP, some serializable schedules are notallowed and conflict detection must be performed on each cache miss.

The final combination of versioning and conflict detection isEager-Optimistic (EO). EO may be a less than optimal choice for HTMsystems: since new transactional versions are written in-place, othertransactions have no choice but to notice conflicts as they occur (i.e.,as cache misses occur). But since EO waits until commit time to detectconflicts, those transactions become “zombies,” continuing to execute,wasting resources, yet are “doomed” to abort.

EO has proven to be useful in STMs and is implemented by Bartok-STM andMcRT. A lazy versioning STM needs to check its write buffer on each readto ensure that it is reading the most recent value. Since the writebuffer is not a hardware structure, this is expensive, hence thepreference for write-in-place eager versioning. Additionally, sincechecking for conflicts is also expensive in an STM, optimistic conflictdetection offers the advantage of performing this operation in bulk.

Contention Management

How a transaction rolls back once the system has decided to abort thattransaction has been described above, but, since a conflict involves twotransactions, the topics of which transaction should abort, how thatabort should be initiated, and when should the aborted transaction beretried need to be explored. These are topics that are addressed byContention Management (CM), a key component of transactional memory.Described below are policies regarding how the systems initiate abortsand the various established methods of managing which transactionsshould abort in a conflict.

Contention Management Policies

A Contention Management (CM) Policy is a mechanism that determines whichtransaction involved in a conflict should abort and when the abortedtransaction should be retried. For example, it is often the case thatretrying an aborted transaction immediately does not lead to the bestperformance. Conversely, employing a back-off mechanism, which delaysthe retrying of an aborted transaction, can yield better performance.STMs first grappled with finding the best contention management policiesand many of the policies outlined below were originally developed forSTMs.

CM Policies draw on a number of measures to make decisions, includingages of the transactions, size of read- and write-sets, the number ofprevious aborts, etc. The combinations of measures to make suchdecisions are endless, but certain combinations are described below,roughly in order of increasing complexity.

To establish some nomenclature, first note that in a conflict there aretwo sides: the attacker and the defender. The attacker is thetransaction requesting access to a shared memory location. Inpessimistic conflict detection, the attacker is the transaction issuingthe load or load exclusive. In optimistic, the attacker is thetransaction attempting to validate. The defender in both cases is thetransaction receiving the attacker's request.

An Aggressive CM Policy immediately and always retries either theattacker or the defender. In LO, Aggressive means that the attackeralways wins, and so Aggressive is sometimes called committer wins. Sucha policy was used for the earliest LO systems. In the case of EP,Aggressive can be either defender wins or attacker wins.

Restarting a conflicting transaction that will immediately experienceanother conflict is bound to waste work—namely interconnect bandwidthrefilling cache misses. A Polite CM Policy employs exponential back-off(but linear could also be used) before restarting conflicts. To preventstarvation, a situation where a process does not have resourcesallocated to it by the scheduler, the exponential back-off greatlyincreases the odds of transaction success after some n retries.

Another approach to conflict resolution is to randomly abort theattacker or defender (a policy called Randomized). Such a policy may becombined with a randomized back-off scheme to avoid unneeded contention.

However, making random choices, when selecting a transaction to abort,can result in aborting transactions that have completed “a lot of work,”which can waste resources. To avoid such waste, the amount of workcompleted on the transaction can be taken into account when determiningwhich transaction to abort. One measure of work could be a transaction'sage. Other methods include Oldest, Bulk TM, Size Matters, Karma, andPolka. Oldest is a simple timestamp method that aborts the youngertransaction in a conflict. Bulk TM uses this scheme. Size Matters islike Oldest but instead of transaction age, the number of read/writtenwords is used as the priority, reverting to Oldest after a fixed numberof aborts. Karma is similar, using the size of the write-set aspriority. Rollback then proceeds after backing off a fixed amount oftime. Aborted transactions keep their priorities after being aborted(hence the name Karma). Polka works like Karma but instead of backingoff a predefined amount of time, it backs off exponentially more eachtime.

Since aborting wastes work, it is logical to argue that stalling anattacker until the defender has finished their transaction would lead tobetter performance. Unfortunately, such a simple scheme easily leads todeadlock.

Deadlock avoidance techniques can be used to solve this problem. Greedyuses two rules to avoid deadlock. The first rule is, if a firsttransaction, T1, has lower priority than a second transaction, T0, or ifT1 is waiting for another transaction, then T1 aborts when conflictingwith T0. The second rule is, if T1 has higher priority than T0 and isnot waiting, then T0 waits until T1 commits, aborts, or starts waiting(in which case the first rule is applied). Greedy provides someguarantees about time bounds for executing a set of transactions. One EPdesign (LogTM) uses a CM policy similar to Greedy to achieve stallingwith conservative deadlock avoidance.

Example MESI coherency rules provide for four possible states in which acache line of a multiprocessor cache system may reside, M, E, S, and I,defined as follows:

Modified (M): The cache line is present only in the current cache, andis dirty; it has been modified from the value in main memory. The cacheis required to write the data back to main memory at some time in thefuture, before permitting any other read of the (no longer valid) mainmemory state. The write-back changes the line to the Exclusive state.

Exclusive (E): The cache line is present only in the current cache, butis clean; it matches main memory. It may be changed to the Shared stateat any time, in response to a read request. Alternatively, it may bechanged to the Modified state when writing to it.

Shared (S): Indicates that this cache line may be stored in other cachesof the machine and is “clean”; it matches the main memory. The line maybe discarded (changed to the Invalid state) at any time.

Invalid (I): Indicates that this cache line is invalid (unused).

TM coherency status indicators (R 132, W 138) may be provided for eachcache line, in addition to, or encoded in the MESI coherency bits. An R132 indicator indicates the current transaction has read from the dataof the cache line, and a W 138 indicator indicates the currenttransaction has written to the data of the cache line.

In another aspect of TM design, a system is designed using transactionalstore buffers. U.S. Pat. No. 6,349,361 titled “Methods and Apparatus forReordering and Renaming Memory References in a Multiprocessor ComputerSystem,” filed Mar. 31, 2000 and incorporated by reference herein in itsentirety, teaches a method for reordering and renaming memory referencesin a multiprocessor computer system having at least a first and a secondprocessor. The first processor has a first private cache and a firstbuffer, and the second processor has a second private cache and a secondbuffer. The method includes the steps of, for each of a plurality ofgated store requests received by the first processor to store a datum,exclusively acquiring a cache line that contains the datum by the firstprivate cache, and storing the datum in the first buffer. Upon the firstbuffer receiving a load request from the first processor to load aparticular datum, the particular datum is provided to the firstprocessor from among the data stored in the first buffer based on anin-order sequence of load and store operations. Upon the first cachereceiving a load request from the second cache for a given datum, anerror condition is indicated and a current state of at least one of theprocessors is reset to an earlier state when the load request for thegiven datum corresponds to the data stored in the first buffer.

The main implementation components of one such transactional memoryfacility are a transaction-backup register file for holdingpre-transaction GR (general register) content, a cache directory totrack the cache lines accessed during the transaction, a store cache tobuffer stores until the transaction ends, and firmware routines toperform various complex functions. In this section a detailedimplementation is described.

IBM zEnterprise EC12 Enterprise Server Embodiment

The IBM zEnterprise EC12 enterprise server introduces transactionalexecution (TX) in transactional memory, and is described in part in apaper, “Transactional Memory Architecture and Implementation for IBMSystem z” of Proceedings Pages 25-36 presented at MICRO-45, 1-5 Dec.2012, Vancouver, British Columbia, Canada, available from IEEE ComputerSociety Conference Publishing Services (CPS), which is incorporated byreference herein in its entirety. “IBM,” “zEnterprise,” “System z,”“EC12,” and/or “MICRO-45” may be subject to trademark rights in variousjurisdictions throughout the world and are used here only in referenceto the products or services properly denominated by the marks to theextent that such trademark rights may exist.).

Table 3 shows an example transaction. Transactions started with TBEGINare not assured to ever successfully complete with TEND, since they canexperience an aborting condition at every attempted execution, e.g., dueto repeating conflicts with other CPUs. This requires that the programsupport a fallback path to perform the same operationnon-transactionally, e.g., by using traditional locking schemes. Thisputs significant burden on the programming and software verificationteams, especially where the fallback path is not automatically generatedby a reliable compiler.

TABLE 3 Example Transaction Code LHI R0,0 *initialize retry count=0 loopTBEGIN *begin transaction JNZ abort *go to abort code if CC1=0 LT R1,lock *load and test the fallback lock JNZ lckbzy *branch if lock busy .. . perform operation . . . TEND *end transaction . . . . . . . . . . .. lckbzy TABORT *abort if lock busy; this *resumes after TBEGIN abort JOfallback *no retry if CC=3 AHI R0, 1 *increment retry count CIJNL R0,6,fallback *give up after 6 attempts PPA R0, TX *random delay based onretry count . . . potentially wait for lock to become free . . . J loop*jump back to retry fallback OBTAIN lock *using Compare&Swap . . .perform operation . . . RELEASE lock . . . . . . . . . . . .

The requirement of providing a fallback path for aborted TransactionExecution (TX) transactions can be onerous. Many transactions operatingon shared data structures are expected to be short, touch only a fewdistinct memory locations, and use simple instructions only. For thosetransactions, the IBM zEnterprise EC12 introduces the concept ofconstrained transactions; under normal conditions, the CPU 114 a (FIG.2) assures that constrained transactions eventually end successfully,albeit without giving a strict limit on the number of necessary retries.A constrained transaction starts with a TBEGINC instruction and endswith a regular TEND. Implementing a task as a constrained ornon-constrained transaction typically results in very comparableperformance, but constrained transactions simplify software developmentby removing the need for a fallback path. IBM's Transactional Executionarchitecture is further described in z/Architecture, Principles ofOperation, Tenth Edition, SA22-7832-09 published September 2012 fromIBM, incorporated by reference herein in its entirety.

A constrained transaction starts with the TBEGINC instruction. Atransaction initiated with TBEGINC must follow a list of programmingconstraints; otherwise the program takes a non-filterableconstraint-violation interruption. Exemplary constraints may include,but not be limited to: the transaction can execute a maximum of 32instructions, all instruction text must be within 256 consecutive bytesof memory; the transaction contains only forward-pointing relativebranches (i.e., no loops or subroutine calls); the transaction canaccess a maximum of 4 aligned octowords (an octoword is 32 bytes) ofmemory; and restriction of the instruction-set to exclude complexinstructions like decimal or floating-point operations. The constraintsare chosen such that many common operations like doubly linkedlist-insert/delete operations can be performed, including the verypowerful concept of atomic compare-and-swap targeting up to 4 alignedoctowords. At the same time, the constraints were chosen conservativelysuch that future CPU implementations can assure transaction successwithout needing to adjust the constraints, since that would otherwiselead to software incompatibility.

TBEGINC mostly behaves like XBEGIN in TSX or TBEGIN on IBM's zEC12servers, except that the floating-point register (FPR) control and theprogram interruption filtering fields do not exist and the controls areconsidered to be zero. On a transaction abort, the instruction addressis set back directly to the TBEGINC instead of to the instruction after,reflecting the immediate retry and absence of an abort path forconstrained transactions.

Nested transactions are not allowed within constrained transactions, butif a TBEGINC occurs within a non-constrained transaction it is treatedas opening a new non-constrained nesting level just like TBEGIN would.This can occur, e.g., if a non-constrained transaction calls asubroutine that uses a constrained transaction internally.

Since interruption filtering is implicitly off, all exceptions during aconstrained transaction lead to an interruption into the operatingsystem (OS). Eventual successful finishing of the transaction relies onthe capability of the OS to page-in the at most 4 pages touched by anyconstrained transaction. The OS must also ensure time-slices long enoughto allow the transaction to complete.

TABLE 4 Transaction Code Example TBEGINC *begin constrained transaction. . . perform operation . . . TEND    *end transaction

Table 4 shows the constrained-transactional implementation of the codein Table 3, assuming that the constrained transactions do not interactwith other locking-based code. No lock testing is shown therefore, butcould be added if constrained transactions and lock-based code weremixed.

When failure occurs repeatedly, software emulation is performed usingmillicode as part of system firmware. Advantageously, constrainedtransactions have desirable properties because of the burden removedfrom programmers.

With reference to FIG. 3, the IBM zEnterprise EC12 processor introducedtransactional execution facility 200. The processor can decode 3instructions per clock cycle; simple instructions are dispatched assingle micro-ops, and more complex instructions are cracked intomultiple micro-ops. Micro-ops (Uops 234 a, 234 b, and 234 c) are writteninto unified issue queue 216, from where they can be issuedout-of-order. Up to two fixed-point, one floating-point, two load/store,and two branch instructions can execute every cycle. Global CompletionTable (GCT) 230 holds every micro-op 234 a, 234 b, and 234 c andtransaction nesting depth (TND) 232. The GCT 230 is written in-order atdecode time, tracks the execution status of each micro-op 234 a, 234 b,and 234 c, and completes instructions when all micro-ops 234 a, 234 b,and 234 c of the oldest instruction group have successfully executed.

Level 1 (L1) data cache 240 is a 96 KB (kilo-byte) 6-way associativecache with 256 byte cache-lines and 4 cycle use latency, coupled to aprivate 1 MB (mega-byte) 8-way associative 2nd-level (L2) data cache 268with 7 cycles use-latency penalty for L1 240 misses. The L1 240 cache isthe cache closest to a processor and Ln cache is a cache at the nthlevel of caching. Both L1 240 and L2 268 caches are store-through. Sixcores on each central processor (CP) chip share a 48 MB 3rd-levelstore-in cache, and six CP chips are connected to an off-chip 384 MB4th-level cache, packaged together on a glass ceramic multi-chip module(MCM). Up to 4 multi-chip modules (MCMs) can be connected to a coherentsymmetric multi-processor (SMP) system with up to 144 cores (not allcores are available to run customer workload).

Coherency is managed with a variant of the MESI protocol. Cache-linescan be owned read-only (shared) or exclusive; the L1 240 and L2 268 arestore-through and thus do not contain dirty lines. The L3 272 and L4caches (not shown) are store-in and track dirty states. Each cache isinclusive of all its connected lower level caches.

Coherency requests are called “cross interrogates” (XI) and are senthierarchically from higher level to lower-level caches, and between theL4s. When one core misses the L1 240 and L2 268 and requests the cacheline from its local L3 272, the L3 272 checks whether it owns the line,and if necessary sends an XI to the currently owning L2 268/L1 240 underthat L3 272 to ensure coherency, before it returns the cache line to therequestor. If the request also misses the L3 272, the L3 272 sends arequest to the L4 (not shown), which enforces coherency by sending XIsto all necessary L3s under that L4, and to the neighboring L4s. Then theL4 responds to the requesting L3 which forwards the response to the L2268/L1 240.

Note that due to the inclusivity rule of the cache hierarchy, sometimescache lines are XI'ed from lower-level caches due to evictions onhigher-level caches caused by associativity overflows from requests toother cache lines. These XIs can be called “LRU XIs”, where LRU standsfor least recently used.

Making reference to yet another type of XI requests, Demote-XIstransition cache-ownership from exclusive into read-only state, andExclusive-XIs transition cache ownership from exclusive into invalidstate. Demote-XIs and Exclusive-XIs need a response back to the XIsender. The target cache can “accept” the XI, or send a “reject”response if it first needs to evict dirty data before accepting the XI.The L1 240/L2 268 caches are store through, but may reject demote-XIsand exclusive XIs if they have stores in their store queues that need tobe sent to L3 before downgrading the exclusive state. A rejected XI willbe repeated by the sender. Read-only-XIs are sent to caches that own theline read-only; no response is needed for such XIs since they cannot berejected. The details of the SMP protocol are similar to those describedfor the IBM z10 by P. Mak, C. Walters, and G. Strait, in “IBM System z10processor cache subsystem microarchitecture”, IBM Journal of Researchand Development, Vol 53:1, 2009, which is incorporated by referenceherein in its entirety.

Transactional Instruction Execution

FIG. 3 depicts example components of an example transactional executionenvironment, including a CPU and caches/components with which itinteracts (such as those depicted in FIGS. 1 and 2). The instructiondecode unit 208 (IDU) keeps track of the current transaction nestingdepth 212 (TND). When the IDU 208 receives a TBEGIN instruction, thenesting depth 212 is incremented, and conversely decremented on TENDinstructions. The nesting depth 212 is written into the GCT 230 forevery dispatched instruction. When a TBEGIN or TEND is decoded on aspeculative path that later gets flushed, the IDU's 208 nesting depth212 is refreshed from the youngest GCT 230 entry that is not flushed.The transactional state is also written into the issue queue 216 forconsumption by the execution units, mostly by the Load/Store Unit (LSU)280, which also has an effective address calculator 236 is included inthe LSU 280. The TBEGIN instruction may specify a transaction diagnosticblock (TDB) for recording status information, should the transactionabort before reaching a TEND instruction.

Similar to the nesting depth, the IDU 208/GCT 230 collaboratively trackthe access register/floating-point register (AR/FPR) modification masksthrough the transaction nest; the IDU 208 can place an abort requestinto the GCT 230 when an AR/FPR-modifying instruction is decoded and themodification mask blocks that. When the instruction becomesnext-to-complete, completion is blocked and the transaction aborts.Other restricted instructions are handled similarly, including TBEGIN ifdecoded while in a constrained transaction, or exceeding the maximumnesting depth.

An outermost TBEGIN is cracked into multiple micro-ops depending on theGR-Save-Mask; each micro-op 234 a, 234 b, and 234 c (including, forexample uop 0, uop 1, and uop2) will be executed by one of the two fixedpoint units (FXUs) 220 to save a pair of GRs 228 into a specialtransaction-backup register file 224, that is used to later restore theGR 228 content in case of a transaction abort. Also the TBEGIN spawnsmicro-ops 234 a, 234 b, and 234 c to perform an accessibility test forthe TDB if one is specified; the address is saved in a special purposeregister for later usage in the abort case. At the decoding of anoutermost TBEGIN, the instruction address and the instruction text ofthe TBEGIN are also saved in special purpose registers for a potentialabort processing later on.

TEND and NTSTG are single micro-op 234 a, 234 b, and 234 c instructions;NTSTG (non-transactional store) is handled like a normal store exceptthat it is marked as non-transactional in the issue queue 216 so thatthe LSU 280 can treat it appropriately. TEND is a no-op at executiontime, the ending of the transaction is performed when TEND completes.

As mentioned, instructions that are within a transaction are marked assuch in the issue queue 216, but otherwise execute mostly unchanged; theLSU 280 performs isolation tracking as described in the next section.

Since decoding is in-order, and since the IDU 208 keeps track of thecurrent transactional state and writes it into the issue queue 216 alongwith every instruction from the transaction, execution of TBEGIN, TEND,and instructions before, within, and after the transaction can beperformed out of order. It is even possible (though unlikely) that TENDis executed first, then the entire transaction, and lastly the TBEGINexecutes. Program order is restored through the GCT 230 at completiontime. The length of transactions is not limited by the size of the GCT230, since general purpose registers (GRs) 228 can be restored from thebackup register file 224.

During execution, the program event recording (PER) events are filteredbased on the Event Suppression Control, and a PER TEND event is detectedif enabled. Similarly, while in transactional mode, a pseudo-randomgenerator may be causing the random aborts as enabled by the TransactionDiagnostics Control.

Tracking for Transactional Isolation

The Load/Store Unit 280 tracks cache lines that were accessed duringtransactional execution, and triggers an abort if an XI from another CPU(or an LRU-XI) conflicts with the footprint. If the conflicting XI is anexclusive or demote XI, the LSU 280 rejects the XI back to the L3 272 inthe hope of finishing the transaction before the L3 272 repeats the XI.This “stiff-arming” is very efficient in highly contended transactions.In order to prevent hangs when two CPUs stiff-arm each other, aXI-reject counter is implemented, which triggers a transaction abortwhen a threshold is met.

The L1 cache directory 240 is traditionally implemented with staticrandom access memories (SRAMs). For the transactional memoryimplementation, the valid bits 244 (64 rows×6 ways) of the directoryhave been moved into normal logic latches, and are supplemented with twomore bits per cache line: the TX-read 248 and TX-dirty 252 bits.

The TX-read 248 bits are reset when a new outermost TBEGIN is decoded(which is interlocked against a prior still pending transaction). TheTX-read 248 bit is set at execution time by every load instruction thatis marked “transactional” in the issue queue. Note that this can lead toover-marking if speculative loads are executed, for example on amispredicted branch path. The alternative of setting the TX-read 248 bitat load completion time was too expensive for silicon area, sincemultiple loads can complete at the same time, requiring many read-portson the load-queue.

Stores execute the same way as in non-transactional mode, but atransaction mark is placed in the store queue (STQ) 260 entry of thestore instruction. At write-back time, when the data from the STQ 260 iswritten into the L1 240, the TX-dirty bit 252 in the L1-directory 256 isset for the written cache line. Store write-back into the L1 240 occursonly after the store instruction has completed, and at most one store iswritten back per cycle. Before completion and write-back, loads canaccess the data from the STQ 260 by means of store-forwarding; afterwrite-back, the CPU 114 a (FIG. 2) can access the speculatively updateddata in the L1 240. If the transaction ends successfully, the TX-dirtybits 252 of all cache-lines are cleared, and also the TX-marks of notyet written stores are cleared in the STQ 260, effectively turning thepending stores into normal stores.

On a transaction abort, all pending transactional stores are invalidatedfrom the STQ 260, even those already completed. All cache lines thatwere modified by the transaction in the L1 240, that is, have theTX-dirty bit 252 on, have their valid bits turned off, effectivelyremoving them from the L1 240 cache instantaneously.

The architecture requires that before completing a new instruction, theisolation of the transaction read- and write-set is maintained. Thisisolation is ensured by stalling instruction completion at appropriatetimes when XIs are pending; speculative out-of order execution isallowed, optimistically assuming that the pending XIs are to differentaddresses and not actually cause a transaction conflict. This designfits very naturally with the XI-vs-completion interlocks that areimplemented on prior systems to ensure the strong memory ordering thatthe architecture requires.

When the L1 240 receives an XI, L1 240 accesses the directory to checkvalidity of the XI'ed address in the L1 240, and if the TX-read bit 248is active on the XI'ed line and the XI is not rejected, the LSU 280triggers an abort. When a cache line with active TX-read bit 248 isLRU'ed from the L1 240, a special LRU-extension vector remembers foreach of the 64 rows of the L1 240 that a TX-read line existed on thatrow. Since no precise address tracking exists for the LRU extensions,any non-rejected XI that hits a valid extension row the LSU 280 triggersan abort. Providing the LRU-extension effectively increases the readfootprint capability from the L1-size to the L2-size and associativity,provided no conflicts with other CPUs 114 a and 114 b (FIGS. 1 and 2)against the non-precise LRU-extension tracking causes aborts.

The store footprint is limited by the store cache size (the store cacheis discussed in more detail below) and thus implicitly by the L2 268size and associativity. No LRU-extension action needs to be performedwhen a TX-dirty 252 cache line is LRU'ed from the L1 240.

Store Cache

In prior systems, since the L1 240 and L2 268 are store-through caches,every store instruction causes an L3 272 store access; with now 6 coresper L3 272 and further improved performance of each core, the store ratefor the L3 272 (and to a lesser extent for the L2 268) becomesproblematic for certain workloads. In order to avoid store queuingdelays, a gathering store cache 264 had to be added, that combinesstores to neighboring addresses before sending them to the L3 272.

For transactional memory performance, it is acceptable to invalidateevery TX-dirty 252 cache line from the L1 240 on transaction aborts,because the L2 268 cache is very close (7 cycles L1 240 miss penalty) tobring back the clean lines. However, it would be unacceptable forperformance (and silicon area for tracking) to have transactional storeswrite the L2 268 before the transaction ends and then invalidate alldirty L2 268 cache lines on abort (or even worse on the shared L3 272).

The two problems of store bandwidth and transactional memory storehandling can both be addressed with the gathering store cache 264. Thecache 264 is a circular queue of 64 entries, each entry holding 128bytes of data with byte-precise valid bits. In non-transactionaloperation, when a store is received from the LSU 280, the store cache264 checks whether an entry exists for the same address, and if sogathers the new store into the existing entry. If no entry exists, a newentry is written into the queue, and if the number of free entries fallsunder a threshold, the oldest entries are written back to the L2 268 andL3 272 caches.

When a new outermost transaction begins, all existing entries in thestore cache are marked closed so that no new stores can be gathered intothem, and eviction of those entries to L2 268 and L3 272 is started.From that point on, the transactional stores coming out of the LSU 280STQ 260 allocate new entries, or gather into existing transactionalentries. The write-back of those stores into L2 268 and L3 272 isblocked, until the transaction ends successfully; at that pointsubsequent (post-transaction) stores can continue to gather intoexisting entries, until the next transaction closes those entries again.

The store cache 264 is queried on every exclusive or demote XI, andcauses an XI reject if the XI compares to any active entry. If the coreis not completing further instructions while continuously rejecting XIs,the transaction is aborted at a certain threshold to avoid hangs.

The LSU 280 requests a transaction abort when the store cache 264overflows. The LSU 280 detects this condition when it tries to send anew store that cannot merge into an existing entry, and the entire storecache 264 is filled with stores from the current transaction. The storecache 264 is managed as a subset of the L2 268: while transactionallydirty lines can be evicted from the L1 240, they have to stay residentin the L2 268 throughout the transaction. The maximum store footprint isthus limited to the store cache size of 64×128 bytes, and it is alsolimited by the associativity of the L2 268. Since the L2 268 is 8-wayassociative and has 512 rows, it is typically large enough to not causetransaction aborts.

If a transaction aborts, the store cache 264 is notified and all entriesholding transactional data are invalidated. The store cache 264 also hasa mark per doubleword (8 bytes) whether the entry was written by a NTSTGinstruction—those doublewords stay valid across transaction aborts.

Millicode-Implemented Functions

Traditionally, IBM mainframe server processors contain a layer offirmware called millicode which performs complex functions like certainCISC instruction executions, interruption handling, systemsynchronization, and RAS. Millicode includes machine dependentinstructions as well as instructions of the instruction set architecture(ISA) that are fetched and executed from memory similarly toinstructions of application programs and the operating system (OS).Firmware resides in a restricted area of main memory that customerprograms cannot access. When hardware detects a situation that needs toinvoke millicode, the instruction fetching unit 204 switches into“millicode mode” and starts fetching at the appropriate location in themillicode memory area. Millicode may be fetched and executed in the sameway as instructions of the instruction set architecture (ISA), and mayinclude ISA instructions.

For transactional memory, millicode is involved in various complexsituations. Every transaction abort invokes a dedicated millicodesubroutine to perform the necessary abort steps. The transaction-abortmillicode starts by reading special-purpose registers (SPRs) holding thehardware internal abort reason, potential exception reasons, and theaborted instruction address, which millicode then uses to store a TDB ifone is specified. The TBEGIN instruction text is loaded from an SPR toobtain the GR-save-mask, which is needed for millicode to know which GRs238 to restore.

The CPU 114 a (FIG. 2) supports a special millicode-only instruction toread out the backup-GRs 224 and copy them into the main GRs 228. TheTBEGIN instruction address is also loaded from an SPR to set the newinstruction address in the PSW to continue execution after the TBEGINonce the millicode abort subroutine finishes. That PSW may later besaved as program-old PSW in case the abort is caused by a non-filteredprogram interruption.

The TABORT instruction may be millicode implemented; when the IDU 208decodes TABORT, it instructs the instruction fetch unit to branch intoTABORT's millicode, from which millicode branches into the common abortsubroutine.

The Extract Transaction Nesting Depth (ETND) instruction may also bemillicoded, since it is not performance critical; millicode loads thecurrent nesting depth out of a special hardware register and places itinto a GR 228. The PPA instruction is millicoded; it performs theoptimal delay based on the current abort count provided by software asan operand to PPA, and also based on other hardware internal state.

For constrained transactions, millicode may keep track of the number ofaborts. The counter is reset to 0 on successful TEND completion, or ifan interruption into the OS occurs (since it is not known if or when theOS will return to the program). Depending on the current abort count,millicode can invoke certain mechanisms to improve the chance of successfor the subsequent transaction retry. The mechanisms involve, forexample, successively increasing random delays between retries, andreducing the amount of speculative execution to avoid encounteringaborts caused by speculative accesses to data that the transaction isnot actually using. As a last resort, millicode can broadcast to otherCPUs other than 114 a which is processing the local transaction, to stopall conflicting work and retry the local transaction before releasingthe other CPUs to continue normal processing. Where multiple CPUs areenabled, their activity must be coordinated to not cause deadlocks, sosome serialization between millicode instances on different CPUs 114 isrequired.

B. Computer Program Product Claim Support

A computer readable storage medium can be a tangible device that canretain and store instructions for use by an instruction executiondevice. The computer readable storage medium may be, for example, but isnot limited to, an electronic storage device, a magnetic storage device,an optical storage device, an electromagnetic storage device, asemiconductor storage device, or any suitable combination of theforegoing. A non-exhaustive list of more specific examples of thecomputer readable storage medium includes the following: a portablecomputer diskette, a hard disk, a random access memory (RAM), aread-only memory (ROM), an erasable programmable read-only memory (EPROMor Flash memory), a static random access memory (SRAM), a portablecompact disc read-only memory (CD-ROM), a digital versatile disk (DVD),a memory stick, a floppy disk, a mechanically encoded device such aspunch-cards or raised structures in a groove having instructionsrecorded thereon, and any suitable combination of the foregoing. Acomputer readable storage medium, as used herein, is not to be construedas being transitory signals per se, such as radio waves or other freelypropagating electromagnetic waves, electromagnetic waves propagatingthrough a waveguide or other transmission media (e.g., light pulsespassing through a fiber-optic cable), or electrical signals transmittedthrough a wire.

Computer readable program instructions described herein can bedownloaded to respective computing/processing devices from a computerreadable storage medium or to an external computer or external storagedevice via a network, for example, the Internet, a local area network, awide area network and/or a wireless network. The network may comprisecopper transmission cables, optical transmission fibers, wirelesstransmission, routers, firewalls, switches, gateway computers and/oredge servers. A network adapter card or network interface in eachcomputing/processing device receives computer readable programinstructions from the network and forwards the computer readable programinstructions for storage in a computer readable storage medium withinthe respective computing/processing device.

Computer readable program instructions for carrying out operations ofthe present invention may be assembler instructions,instruction-set-architecture (ISA) instructions, machine instructions,machine dependent instructions, microcode, firmware instructions,state-setting data, or either source code or object code written in anycombination of one or more programming languages, including an objectoriented programming language such as Smalltalk, C++ or the like, andconventional procedural programming languages, such as the “C”programming language or similar programming languages. The computerreadable program instructions may execute entirely on the user'scomputer, partly on the user's computer, as a stand-alone softwarepackage, partly on the user's computer and partly on a remote computeror entirely on the remote computer or server. In the latter scenario,the remote computer may be connected to the user's computer through anytype of network, including a local area network (LAN) or a wide areanetwork (WAN), or the connection may be made to an external computer(for example, through the Internet using an Internet Service Provider).In some embodiments, electronic circuitry including, for example,programmable logic circuitry, field-programmable gate arrays (FPGA), orprogrammable logic arrays (PLA) may execute the computer readableprogram instructions by utilizing state information of the computerreadable program instructions to personalize the electronic circuitry,in order to perform aspects of the present invention.

Aspects of the present invention are described herein with reference toflowchart illustrations and/or block diagrams of methods, apparatus(systems), and computer program products according to embodiments of theinvention. It will be understood that each block of the flowchartillustrations and/or block diagrams, and combinations of blocks in theflowchart illustrations and/or block diagrams, can be implemented bycomputer readable program instructions.

These computer readable program instructions may be provided to aprocessor of a general purpose computer, special purpose computer, orother programmable data processing apparatus to produce a machine, suchthat the instructions, which execute via the processor of the computeror other programmable data processing apparatus, create means forimplementing the functions/acts specified in the flowchart and/or blockdiagram block or blocks. These computer readable program instructionsmay also be stored in a computer readable storage medium that can directa computer, a programmable data processing apparatus, and/or otherdevices to function in a particular manner, such that the computerreadable storage medium having instructions stored therein comprises anarticle of manufacture including instructions which implement aspects ofthe function/act specified in the flowchart and/or block diagram blockor blocks.

The computer readable program instructions may also be loaded onto acomputer, other programmable data processing apparatus, or other deviceto cause a series of operational steps to be performed on the computer,other programmable apparatus or other device to produce a computerimplemented process, such that the instructions which execute on thecomputer, other programmable apparatus, or other device implement thefunctions/acts specified in the flowchart and/or block diagram block orblocks.

The flowchart and block diagrams in the Figures illustrate thearchitecture, functionality, and operation of possible implementationsof systems, methods, and computer program products according to variousembodiments of the present invention. In this regard, each block in theflowchart or block diagrams may represent a module, segment, or portionof instructions, which comprises one or more executable instructions forimplementing the specified logical function(s). In some alternativeimplementations, the functions noted in the block may occur out of theorder noted in the figures. For example, two blocks shown in successionmay, in fact, be executed substantially concurrently, or the blocks maysometimes be executed in the reverse order, depending upon thefunctionality involved. It will also be noted that each block of theblock diagrams and/or flowchart illustration, and combinations of blocksin the block diagrams and/or flowchart illustration, can be implementedby special purpose hardware-based systems that perform the specifiedfunctions or acts or carry out combinations of special purpose hardwareand computer instructions.

C. Computer Program Product Claim Support

An embodiment of a possible hardware and software environment forsoftware and/or methods according to the present invention will now bedescribed in detail with reference to the Figures. FIG. 4 depicts ablock diagram of components of a computing device 400, in accordancewith an illustrative embodiment of the present invention. It should beappreciated that FIG. 4 provides only an illustration of oneimplementation and does not imply any limitations with regard to theenvironments in which different embodiments may be implemented. Manymodifications to the depicted environment may be made. It should beappreciated FIG. 4 provides only an illustration of one implementationand does not imply any limitations with regard to the environments inwhich different embodiments may be implemented.

The computing environment of FIG. 4 is, in many respects, representativeof the various computer subsystem(s) in the present invention.Accordingly, several portions of the computing environment will now bediscussed in the following paragraphs.

Computing device 400 includes communications fabric 402, which providescommunications between computer processor(s) 404, memory 406, persistentstorage 408, communications unit 410, and input/output (I/O)interface(s) 412. Communications fabric 402 can be implemented with anyarchitecture designed for passing data and/or control informationbetween processors (such as microprocessors, communications and networkprocessors, etc.), system memory, peripheral devices, and any additionalhardware components within a system. For example, communications fabric402 can be implemented with one or more buses.

Computing device 400 is capable of communicating with other computersubsystems via network 401. Network 401 can be, for example, a localarea network (LAN), a wide area network (WAN) such as the Internet, or acombination of the two, and can include wired, wireless, or fiber opticconnections. In general, network 401 can be any combination ofconnections and protocols that will support communications betweencomputing device 400 and other computing devices.

Memory 406 and persistent storage 408 are computer-readable storagemedia. In one embodiment, memory 406 includes random access memory (RAM)and cache memory 414. In general, memory 406 can include any suitablevolatile or non-volatile computer-readable storage media.

In some embodiments, regulation logic 420 may be stored for execution byone or more of the respective computer processors 404 of computingdevice 400 via one or more memories of memory 406 of computing device400. In the depicted embodiment, persistent storage 408 includes amagnetic hard disk drive. Alternatively, or in addition to a magnetichard disk drive, persistent storage 408 can include a solid state harddrive, a semiconductor storage device, read-only memory (ROM), erasableprogrammable read-only memory (EPROM), flash memory, or any othercomputer-readable storage media that is capable of storing programinstructions or digital information. In some embodiments, regulationlogic 420 may be implemented using logic gates.

The media used by persistent storage 408 may also be removable. Forexample, a removable hard drive may be used for persistent storage 408.Other examples include optical and magnetic disks, thumb drives, andsmart cards that are inserted into a drive for transfer onto anothercomputer-readable storage medium that is also part of persistent storage408.

Communications unit 410, in the examples, provides for communicationswith other data processing systems or devices, including computingdevice 400. In the examples, communications unit 410 includes one ormore network interface cards. Communications unit 410 may providecommunications through the use of either or both physical and wirelesscommunications links.

I/O interface(s) 412 allows for input and output of data with otherdevices that may be connected to computing device 400. For example, I/Ointerface 412 may provide a connection to external devices 416 such as akeyboard, keypad, camera, a touch screen, and/or some other suitableinput device. External devices 416 can also include portablecomputer-readable storage media such as, for example, thumb drives,portable optical or magnetic disks, and memory cards. Software and dataused to practice embodiments of the present invention, e.g., regulationlogic 420 can be stored on such portable computer-readable storage mediaand can be loaded onto persistent storage 408 of computing device 400via I/O interface(s) 412 of computing device 400. It should be notedthat, in some embodiments, regulation logic 420 is implemented as ahardware module.

Display 418 provides a mechanism to display data to a user and may be,for example, a computer monitor.

The logic described herein is identified based upon the application forwhich it is implemented in a specific embodiment of the invention.However, it should be appreciated that any particular logic nomenclatureherein is used merely for convenience, and thus the invention should notbe limited to use solely in any specific application identified and/orimplied by such nomenclature.

Regulation logic 420 detects the cause of the termination of thetransaction and selects the parameters for the retrying of theoperation. A transaction is a group of instructions that operate in anatomic manner on a data structure in memory, as viewed by other CPUs andthe I/O subsystem. For a transaction to be complete the changes need tobe finalized and made permanent in their entirety. The processing of atransaction can either be successful or fail, it cannot be partiallycompleted. In additional embodiments, a transaction is an individual, orindivisible, operation which is part of a larger operation. One eventwhich can cause the termination of the transaction before completion isa transient condition. A transient condition, or transient property ofan element of the system, is one which is temporary. Transientconditions may be, for example asynchronous interruptions, another CPUtrying to access memory used within the transaction, or another threadon the same CPU causing a cache line to be evicted using an algorithm ofleast recently used (LRU). Regulation logic 420 detects the transientcondition or element, or another cause of the transaction beingterminated. In additional embodiments, regulation logic 420 records thecause of the transaction being terminated in a repository, such as, forexample, memory 406, persistent storage 408, or as an internal hardwarelogic state.

Regulation logic 420 also controls the procedure performed by computingdevice 400 once regulation logic 420 determines the cause of thepremature termination of the transaction. Regulation logic 420 decidesif the cause of the premature termination of the transaction can befixed with a series of retries of the transaction, or other methods ofallowing the transaction more attempts to be completed. In oneembodiment, regulation logic 420 permits the transaction a predeterminednumber of retries to be completed. The predetermined number of retriesdoes not guarantee a successful transaction, but can be used forassistance in future transactions to increase the speed of thetransaction or to anticipate failures.

In additional embodiments, regulation logic 420 can permit thetransaction to retry until the transaction is successful. In a portionof the additional embodiments, regulation logic 420 may recordinformation related to the transaction, this information can be, forexample, the cause of the premature termination, the number of retries,and the successful transaction. This information can potentially be usedin future transactions which fail to find a known solution and decreasethe time for future transaction to be successful. The programs describedherein are identified based upon the application for which they areimplemented in a specific embodiment of the invention. However, itshould be appreciated that any particular logic nomenclature herein isused merely for convenience, and thus the invention should not belimited to use solely in any specific application identified and/orimplied by such nomenclature.

II. Example Embodiment

FIG. 5 shows flowchart 500 depicting an approach according to thepresent invention. FIG. 6 shows regulation logic 420 for performing atleast some of the steps of flowchart 500. This approach will now bediscussed, over the course of the following paragraphs, with extensivereference to FIG. 5 and FIG. 6.

Processing begins at step 502, where transaction module (“mod”) 602receives a begin-transaction instruction (e.g., T_BEGIN, TXBEGIN, orXBEGIN among others) indicating the beginning of a transaction. Thetransaction comprises the begin-transaction instruction and all theinstructions immediately following the begin-transaction programinstruction up to, and including an end-transaction instruction (e.g.,T_END, TXEND, or XEND among others) that corresponds to thebegin-transaction instruction.

Processing proceeds to step 504, where initialization mod 604 carriesout initialization operations necessary to enable processing of thetransaction to begin. Initialization operations may include, but are notlimited to: (i) setting a transaction indicator to indicate atransaction is currently being processed; (ii) setting a metric counterto zero; (iii) preparing for a rollback situation by keeping a backupcopy of the processor state at the beginning of transaction processing;and/or (iv) invoking flag processing mod 608 to indicate theend-of-transaction has not been detected.

A transaction indicator may be implemented in software, hardware, or acombination of the two. In some embodiments of the present invention,the transaction indicator is a software implementation using a Booleanflag, wherein a zero value indicates that no transaction is beingprocessed, and a one value indicates that a transaction is beingprocessed. Additionally, in some embodiments of the present invention,the transaction indicator is a hardware implementation and a statusregister is used to indicate whether or not a transaction is beingprocessed. Further, in some embodiments of the present invention, nestedtransactions are supported, and the transaction indicator is a counterthat is incremented each time a begin-transaction is encountered, anddecremented each time an end-transaction is encountered, thus,indicating all nested transactions are complete when the value in thetransaction indicator reaches zero.

A metric counter may be used by the instruction processor to determinehow far the current instruction is from the end of the transaction(e.g., how many instructions remain in the transaction). When abegin-transaction instruction is encountered, a metric countercorresponding to the current transaction is initialized to zero. In someembodiments of the present invention, the metric counter is incrementedonce for each instruction identified during a speculative look aheadoperation, and the metric counter is frozen when an end-transactioninstruction is encountered during the speculative look ahead operation(i.e., the metric counter will contain the total number of instructionsincluded in the transaction corresponding to the metric counter). Insome embodiments, the metric counter is decremented for each completedinstruction within the transaction.

In the event that a transaction is unable to successfully complete, theprocessor shall be able to perform a rollback (e.g., restore) of theenvironment (e.g., transaction memory, registers, variables, and thelike) to a state corresponding to the environment at the beginning ofthe transaction operation. In some embodiments of the present invention,a rollback is necessary when another process causes interference byattempting to access the transactional memory corresponding to thetransaction. If the interference causes the transaction to halt (alsoreferred to herein as an abort of the transaction) without reaching theend-transaction instruction, all processing performed during thetransaction is discarded and a rollback operation is performed torestore the environment to the state equivalent to that at the start ofthe transaction.

When an environment utilizes speculative look ahead (i.e., when theenvironment fetches and decodes program instructions prior toexecution), an indicator (e.g., TXEND_INSIGHT) may be maintained todetermine whether an end-of-transaction has been detected during thespeculative look ahead operation. When a begin-transaction instructionis processed, the indicator (e.g., TXEND_INSIGHT) may be initialized toindicate that no end-transaction has been detected.

Processing proceeds to step 506, where flag processing mod 608 receivesa request to update an indicator (e.g., a flag) to contain a specifiedvalue. The received request may identify a flag and the operation thatis to be performed on the flag. In some embodiments of the presentinvention, the flag is TXEND_INSIGHT, and the operation is indicating noend-transaction instruction has been detected. Additionally, in someembodiments of the present invention, additional operations are carriedout, such as, when TXEND_INSIGHT is changed to indicate that noend-transaction instruction has been detected, the metric counter isunfrozen to enable the metric counter to be incremented while thespeculative look ahead operation proceeds.

Processing proceeds to step 508, where monitor mod 610 monitors thespeculative look ahead operation and detect when an end-transactioninstruction is encountered. Speculative look ahead is an operation thatpredicts an execution path that is likely to be followed in the future,however, it is possible that the prediction is incorrect and thepredicted execution path is not actually followed. Predicting anexecution path that is not actually followed may occur when thespeculative look ahead operation encounters a branch instruction in theinstruction stream. Since the information that determines the actualbehavior of the branch instruction may not yet be available, thespeculative look ahead operation may predict the behavior of the branchinstruction based on previous behavior (e.g., including but not limitedto local and global branch history, branch address history, branchhistory vectors and/or a return address stack), static branch predictioninformation encoded in the instruction, branch policies tangiblyincorporated into the processor (e.g., backward branches, i.e., branchesto lower addresses than the address of a branch instruction may bepredicted as “taken”), meta data, or the like.

Instead of waiting for the information that determines the actualbehavior of the branch instruction to become available, the speculativelook ahead operation proceeds on a predicted execution path. In someembodiments of the present invention, a backup copy of the metriccounter is maintained each time a branch instruction is encountered. Ifit is later determined that an incorrect execution path was predictedfor the branch, the metric counter can be restored to the value it waswhen the branch instruction was being speculatively looked at, and thespeculative look ahead operation can resume following a differentexecution path. In some embodiments of the present invention,instruction decode is also performed speculatively ahead of instructionexecution and at least one decoupling queue (e.g., dispatch queue, issuequeue, reorder queue, and the like) is available to store speculativelydecoded instructions after the instruction decode operation and prior toinstruction execution.

If the speculative look ahead operation encounters an end-transactioninstruction, monitor mod 610 performs tasks to indicate that anend-transaction has been encountered (i.e., an end-transaction has beendetected). Tasks to be performed include, but are not limited to: (i)freezing the metric counter; and/or (ii) notifying flag processing mod608 that that an indicator (e.g., a flag) such as TXEND_INSIGHT is to beset to indicate an end-transaction instruction has been encountered.

If an end-transaction instruction has been identified, and it turns outthat the predicted execution path is incorrect, then the effects ofdetecting the end-transaction instruction shall be rolled back to thepoint at which the incorrect branch prediction was determined. Therollback operation may include restoring the speculative look aheadoperation to the point at which the incorrect branch prediction wasdetermined. The tasks included in the rollback operation may include,but are not limited to: (i) unfreezing the metric counter; (ii)restoring (e.g., rolling back) the metric counter to the state (i.e.,value) of the metric counter when the speculative look ahead operationwas looking at the branch instruction; and/or (iii) notifying flagprocessing mod 608 that that an indicator (e.g., a flag) such asTXEND_INSIGHT is to be set to indicate that no end-transactioninstruction has been encountered.

Processing proceeds to step 510, where flag processing mod 608 receivesa request to update an indicator (e.g., a flag) to contain a specifiedvalue. The received request may identify a flag and an operation that isto be performed on the flag. In some embodiments of the presentinvention, the flag is TXEND_INSIGHT, and the operation is indicatingthat an end-transaction instruction has been detected. Additionally, insome embodiments of the present invention, additional operations arecarried out, such as when TXEND_INSIGHT is altered to indicate that anend-transaction instruction has been detected, the metric counter isfrozen to allow calculations to determine how close the currentinstruction is to the end-transaction instruction.

Processing proceeds to step 512, where interference mod 612 receives anotification identifying another process that is attempting to accesstransactional memory corresponding to an active transaction. Inaccordance with one embodiment of the present invention, interference isdetected in conjunction with the tracking of read and write sets ofmemory that has been the subject of accesses of the present transaction,and further explained with reference to FIGS. 1, 2, and 3 herein.Interference mod 612 may analyze the circumstances corresponding to theinterference. The analysis may include, but is not limited to: (i)determining if delaying the requested halt of the transaction is anoption (it should be noted that the terms “halt” and “abort” are usedinterchangeably herein and have the same meaning); (ii) determining ifthe transaction is near the end (i.e., an end-transaction instructionhas been encountered by the speculative look ahead operation); and/or(iii) if both processes are transactions, determining which transactionis closer to completion.

Processing ends with step 514, where determination mod 614 determines anappropriate action to take with regard to the current transaction.Possible actions may include, but are not limited to: (i) delay the haltrequest, and continue processing the transaction; (ii) halt thetransaction and perform a rollback operation; and/or (iii) request thatthe interfering transaction halt execution.

Delaying the halt request may include the interfering process waitingfor the transaction to complete before the interfering process canobtain the requested data. In this example embodiment of the presentinvention: (i) a transaction is being processed; (ii) an interferenceoccurs; and (iii) the speculative look ahead operation has encounteredan end-transaction instruction. Determination mod 614 determines thatthe transaction will continue to be executed, however, it is laterdetermined that the encountered end-transaction instruction was in anincorrect execution path, and therefore no end-transaction instructionhas been encountered, resulting in the transaction being halted.

When a transaction halts, no additional instructions corresponding tothe transaction are run, and a rollback operation may be performed sothe processing environment appears as if the transaction had never begunprocessing. In some embodiments of the present invention, whenever atransaction ends (e.g., halts or runs to completion), a transactionindicator is updated to indicate the transaction is no longer beingprocessed, freeing the processor from any limitations in place duringtransaction processing.

In some embodiments of the current invention, the interferingtransaction is of a lower priority, and the interfering transactionhalts, allowing the current transaction to continue processing. In someembodiments of the present invention, the interfering transaction isoperating with a higher priority, and therefore the current transactionhalts. Additionally, in some embodiments of the present invention, noend-transaction instruction has been encountered by the speculative lookahead operation, and therefore the current transaction halts.

As noted above, interference mod 612 receives access requests fromremote processes and determines interference. When no interference isdetected, and a local processor includes the requested data, a responsewith the requested data is provided. In some embodiments of the presentinvention, the interference module provides the requested data as wellas an indication that a present transaction of the present processor isto be aborted (action (ii), discussed above).

In some embodiments of the present invention, a decision is made as towhether to provide the data immediately (and cause a transaction abortresponsive to such determination), or to defer a response (action (i),above). When a response is deferred, at a later time, responsive tocompleting the present transaction interference mod 612 provides thedata corresponding to a deferred response. Further, when the possibilityof a deadlock has been detected, and a present transaction is to beaborted, interference module is notified to provide the data inconjunction with the initiation of a transaction abort.

It should be noted as described above that when an abort operation isdelayed, or held, both the instruction to abort as well as the data ofthe interfering transaction are withheld from processing.

III. Further Comments and/or Embodiments

Some embodiments of the present invention recognize the following facts,potential problems, and/or potential areas for improvement with respectto the current state of the art: (i) when deciding to make decisionsabout interference between at least two transactions, and deciding whichinstruction to abort, it is desirable to know which transaction will beending shortly (e.g., whether holding off on responding to a transactionmay prevent an interference); (ii) when an interference must cause onetransaction to abort, it may be desirable to halt the transaction thatis not close to completing (e.g., do not sacrifice a transaction thatwould only require a few instruction to be successfully completed);and/or (iii) it is desirable to offer a compatible way of indicating theimpending end of a transaction that allows computer code and computerarchitecture to remain backwardly compatible.

FIG. 7 depicts a flowchart of the operational steps of regulation logic420, within the computing environment of FIG. 4, in accordance with oneembodiment of the present invention. Flowchart 700 depicts the stepstaken by regulation logic 420 to control the steps taken by the hardwareor software. It should be appreciated that FIG. 7 provides only anillustration of one implementation and does not imply any limitationswith regard to the environments in which different embodiments may beimplemented. Many modifications to the depicted environment may be made.

In step 702, regulation logic 420 detects transaction execution. Thisstep is carried out in the process described above in FIG. 5.

In step 704, regulation logic 420 performs TX initialization. This stepis described above in reference to FIG. 2 and FIG. 5.

In step 706, regulation logic 420 executes the transaction. This step iscarried out in the process described above in FIG. 5.

In decision 708, regulation logic 420 determines if the transaction wasaborted. Regulation logic 420 detects the transaction has aborted asdescribed above. An abort may be caused by, for example, a haltoperation, also referred to as an abort operation, which, as usedherein, refers to an operation responsive to a condition where twotransactions have been detected to interfere and where at least onetransaction must be aborted and the state of the processor is reset tothe state at the beginning of the aborted transaction by performing arollback. If regulation logic 420 determines the transaction was aborted(yes branch, proceed to decision 710), regulation logic, 420 determinesif the cause of the abort was due to a transient condition. Ifregulation logic 420 determines the transaction was not aborted (nobranch, proceed to step 712), regulation logic 420 completes thetransaction.

In decision 710, regulation logic 420 determines if the cause of theabort was due to a transient condition. A transient condition is acondition which may show up during one invocation of a transaction butnot in a following transaction. Such conditions may be, for example,interference from another CPU, an asynchronous interruption, anotherthread LRUing an entry out of the L1 cache with TX_read or TX_dirty bitsset. Non-transient conditions that may cause an abortion may include,for example, issuing an illegal instruction or, when running singlethreaded, exceeding a transaction footprint. If regulation logic 420determines the cause of the termination is due to a transient condition(yes branch, proceed to step 716), regulation logic 420 selects theretry parameters. If regulation logic 420 determines the cause of thetermination is not due to a transient condition (no branch, proceed tostep 714), regulation logic 420 aborts the operation.

In step 712, regulation logic 420 completes the transaction. Regulationlogic 420 completes the execution of the transaction till thetransaction has been completed.

In step 714, regulation logic 420 aborts the operation. Regulation logic420 aborts the operation because the cause of the error is beyond atransient condition or element, or the transaction cannot be completedfor other reasons. This can be due to, for example, a loss ofinformation necessary for the transaction, an invalid instruction, afootprint overflow, a transaction taking too long to execute, a hardwarefailure, or other mechanical or technical issues which can arise whichwould not be related to a transient condition or element.

In step 716, regulation logic 420 selects the retry parameters. Thisstep selects what actions regulation logic 420 permits the hardware ofFIG. 4 to perform to attempt to resolve the transient condition. FIG. 8explains the step of selecting the retry parameters in greater detail.

In step 718, regulation logic 420 performs the permitted action.Regulation logic 420 performs the actions approved, and selected in theprocess performed in FIG. 7. Regulation logic 420 may, for example,automatically perform these actions once the permitted actions areselected, or regulation logic 420 may wait a predetermined time beforestarting the permitted action. In additional embodiments, regulationlogic 420 may permit numerous permitted actions.

In step 720, regulation logic 420 resets to initial transaction state.In some embodiments, regulation logic 420 determines if the permittedaction allowed the transaction to succeed. Regulation logic 420determines if the transaction succeeded by detecting if the transactionwas able to be performed fully to completion, without interruption orissues. If regulation logic 420 determines the permitted action allowedthe transaction to succeed (yes branch, proceed to END), regulationlogic 420 concludes the operation. If regulation logic 420 determinesthe permitted action causes the transaction to fail (no branch, proceedto 204), regulation logic 420 reselects the retry parameters. In anattempt to retry the transaction with a new set of parameters to attemptto allow the transaction to succeed.

FIG. 8 depicts a flowchart of the operational steps of regulation logic420 operating within the computing environment of FIG. 1, in accordancewith one embodiment of the present invention. Flowchart 800 depicts thesteps taken by regulation logic 420 to structure the limitations of theactions performed by the computing device. It should be appreciated thatFIG. 8 provides only an illustration of one implementation and does notimply any limitations with regard to the environments in which differentembodiments may be implemented.

In decision 802, regulation logic 420 determines if retry is permitted.Regulation logic 420 may, in some instances, retry the transaction inorder to resolve the conflict. The retry of the transaction can be, forexample, because regulation logic 420 has solved the conflict before,the transient condition is down to regulation logic 420, or retrying thetransaction can potentially fix the transient condition. If regulationlogic 420 determines that the retry is permitted (yes branch, proceed tostep 804), regulation logic 420 sets the parameters for the retry. Ifregulation logic 420 determines the retry is not permitted (no branch,proceed to step 714), regulation logic 420 aborts the operation.

In step 804, regulation logic 420 sets parameter. Regulation logic 420sets the parameter which computing device 400 performs in an attempt tohave the transaction succeed. These parameters can be, for example,number of retries which computing device 400 performs in an attempt toallow the transaction to succeed, time limit which computing device 400is allowed to perform retries, or the action which is performed by eachretry. In additional embodiments there is a predetermined delay inresponding to coherency requests. In additional embodiments, regulationlogic 420 retries the transaction in the same manner as the firstinstance which failed, or regulation logic 420 retries the transactionin a different manner than the first instance which failed. Inadditional embodiments, regulation logic 420 may set the parameters torecord each retry, the actions performed by computing device 400, or ifthe retry was successful or a failure.

In decision 806, regulation logic 420 determines if occurrence isrecorded. In certain instances where a transaction is terminated and itis known that the termination is caused by a transient condition,regulation logic 420 records information regarding the instance, whichcan be, for example, the number of retries, the transient condition, ifthe retries solved the transient condition, or the solution to thetransient condition for future use. If regulation logic 420 determinesthe information regarding the occurrence is to be recorded (yes branch,proceed to step 718), regulation logic 420 records the informationregarding the occurrence (see step 808). If regulation logic 420determines the information regarding the occurrence is to not berecorded (no branch, proceed to step 714), regulation logic 420 proceedsto step 718 (see FIG. 7).

The present invention may be a system, a method, and/or a computerprogram product. The computer program product may include a computerreadable storage medium (or media) having computer readable programinstructions thereon for causing a processor to carry out aspects of thepresent invention.

The computer readable storage medium can be a tangible device that canretain and store instructions for use by an instruction executiondevice. The computer readable storage medium may be, for example, but isnot limited to, an electronic storage device, a magnetic storage device,an optical storage device, an electromagnetic storage device, asemiconductor storage device, or any suitable combination of theforegoing. A non-exhaustive list of more specific examples of thecomputer readable storage medium includes the following: a portablecomputer diskette, a hard disk, a random access memory (RAM), aread-only memory (ROM), an erasable programmable read-only memory (EPROMor Flash memory), a static random access memory (SRAM), a portablecompact disc read-only memory (CD-ROM), a digital versatile disk (DVD),a memory stick, a floppy disk, a mechanically encoded device such aspunch-cards or raised structures in a groove having instructionsrecorded thereon, and any suitable combination of the foregoing. Acomputer readable storage medium, as used herein, is not to be construedas being transitory signals per se, such as radio waves or other freelypropagating electromagnetic waves, electromagnetic waves propagatingthrough a waveguide or other transmission media (e.g., light pulsespassing through a fiber-optic cable), or electrical signals transmittedthrough a wire.

Computer readable program instructions described herein can bedownloaded to respective computing/processing devices from a computerreadable storage medium or to an external computer or external storagedevice via a network, for example, the Internet, a local area network, awide area network and/or a wireless network. The network may comprisecopper transmission cables, optical transmission fibers, wirelesstransmission, routers, firewalls, switches, gateway computers and/oredge servers. A network adapter card or network interface in eachcomputing/processing device receives computer readable programinstructions from the network and forwards the computer readable programinstructions for storage in a computer readable storage medium withinthe respective computing/processing device.

Computer readable program instructions for carrying out operations ofthe present invention may be assembler instructions,instruction-set-architecture (ISA) instructions, machine instructions,machine dependent instructions, microcode, firmware instructions,state-setting data, or either source code or object code written in anycombination of one or more programming languages, including an objectoriented programming language such as Smalltalk, C++ or the like, andconventional procedural programming languages, such as the “C”programming language or similar programming languages. The computerreadable program instructions may execute entirely on the user'scomputer, partly on the user's computer, as a stand-alone softwarepackage, partly on the user's computer and partly on a remote computeror entirely on the remote computer or server. In the latter scenario,the remote computer may be connected to the user's computer through anytype of network, including a local area network (LAN) or a wide areanetwork (WAN), or the connection may be made to an external computer(for example, through the Internet using an Internet Service Provider).In some embodiments, electronic circuitry including, for example,programmable logic circuitry, field-programmable gate arrays (FPGA), orprogrammable logic arrays (PLA) may execute the computer readableprogram instructions by utilizing state information of the computerreadable program instructions to personalize the electronic circuitry,to perform aspects of the present invention.

Aspects of the present invention are described herein with reference toflowchart illustrations and/or block diagrams of methods, apparatus(systems), and computer program products according to embodiments of theinvention. It will be understood that each block of the flowchartillustrations and/or block diagrams, and combinations of blocks in theflowchart illustrations and/or block diagrams, can be implemented bycomputer readable program instructions.

These computer readable program instructions may be provided to aprocessor of a general purpose computer, special purpose computer, orother programmable data processing apparatus to produce a machine, suchthat the instructions, which execute via the processor of the computeror other programmable data processing apparatus, create means forimplementing the functions/acts specified in the flowchart and/or blockdiagram block or blocks. These computer readable program instructionsmay also be stored in a computer readable storage medium that can directa computer, a programmable data processing apparatus, and/or otherdevices to function in a particular manner, such that the computerreadable storage medium having instructions stored therein comprises anarticle of manufacture including instructions which implement aspects ofthe function/act specified in the flowchart and/or block diagram blockor blocks.

The computer readable program instructions may also be loaded onto acomputer, other programmable data processing apparatus, or other deviceto cause a series of operational steps to be performed on the computer,other programmable apparatus or other device to produce a computerimplemented process, such that the instructions which execute on thecomputer, other programmable apparatus, or other device implement thefunctions/acts specified in the flowchart and/or block diagram block orblocks.

The flowchart and block diagrams in the Figures illustrate thearchitecture, functionality, and operation of possible implementationsof systems, methods, and computer program products according to variousembodiments of the present invention. In this regard, each block in theflowchart or block diagrams may represent a module, segment, or portionof instructions, which comprises one or more executable instructions forimplementing the specified logical function(s). In some alternativeimplementations, the functions noted in the block may occur out of theorder noted in the figures. For example, two blocks shown in successionmay, in fact, be executed substantially concurrently, or the blocks maysometimes be executed in the reverse order, depending upon thefunctionality involved. It will also be noted that each block of theblock diagrams and/or flowchart illustration, and combinations of blocksin the block diagrams and/or flowchart illustration, can be implementedby special purpose hardware-based systems that perform the specifiedfunctions or acts or carry out combinations of special purpose hardwareand computer instructions.

IV. Definitions

Present invention: should not be taken as an absolute indication thatthe subject matter described by the term “present invention” is coveredby either the claims as they are filed, or by the claims that mayeventually issue after patent prosecution; while the term “presentinvention” is used to help the reader to get a general feel for whichdisclosures herein that are believed as maybe being new, thisunderstanding, as indicated by use of the term “present invention,” istentative and provisional and subject to change over the course ofpatent prosecution as relevant information is developed and as theclaims are potentially amended.

Embodiment: see definition of “present invention” above—similar cautionsapply to the term “embodiment.”

and/or: inclusive or; for example, A, B “and/or” C means that at leastone of A or B or C is true and applicable.

User/subscriber: includes, but is not necessarily limited to, thefollowing: (i) a single individual human; (ii) an artificialintelligence entity with sufficient intelligence to act as a user orsubscriber; and/or (iii) a group of related users or subscribers.

Module/Sub-Module: any set of hardware, firmware and/or software thatoperatively works to do some kind of function, without regard to whetherthe module is: (i) in a single local proximity; (ii) distributed over awide area; (iii) in a single proximity within a larger piece of softwarecode; (iv) located within a single piece of software code; (v) locatedin a single storage device, memory or medium; (vi) mechanicallyconnected; (vii) electrically connected; and/or (viii) connected in datacommunication.

Computer: any device with significant data processing and/or machinereadable instruction reading capabilities including, but not limited to:desktop computers, mainframe computers, laptop computers,field-programmable gate array (FPGA) based devices, smart phones,personal digital assistants (PDAs), body-mounted or inserted computers,embedded device style computers, application-specific integrated circuit(ASIC) based devices.

What is claimed is:
 1. A method for resolving terminated transactions ina transactional memory environment, the method comprising: initiating ahardware transaction in a computing environment, wherein the hardwaretransaction accesses a memory location, and wherein the hardwaretransaction includes a transaction begin indicator and a transaction endindicator; detecting a conflicting access of the memory location whileexecuting the hardware transaction, wherein: the conflicting access ofthe memory location is an intervening store to the memory location in anear-end transaction processing mode; and the near-end transactionprocessing mode indicates that the conflicting access is near end ofcompletion based on, at least the transaction end indicator; abortingthe hardware transaction based on the conflicting access of the memorylocation and the near-end transaction processing mode of the conflictingaccess of the memory location; determining, by hardware, that theconflicting access of the memory location is a transient condition; andreinitiating the hardware transaction, wherein reinitiating the hardwaretransaction occurs only in response to determining that the conflictingaccess of the memory location is a transient condition.
 2. The method ofclaim 1, further comprising: generating, by software, a retry parameterbased on, at least, the conflicting access of the memory, wherein theretry parameter indicates, to hardware, a maximum number of attempts toreinitiate the hardware transaction prior to aborting the hardwaretransaction.
 3. The method of claim 2, further comprising: determining,by hardware, that the retry parameter has not been met; and whereinreinitiating the hardware transaction is based on the determination thatthe retry parameter has not been met.
 4. The method of claim 1, whereinthe transient condition is a condition which may cause conflictingaccess of the memory location during a first invocation of the hardwaretransaction, but which may not result in conflicting access of thememory location during a subsequent invocation of the hardwaretransaction.
 5. The method of claim 1, wherein: the step of initiatingthe hardware transaction in the computing environment is performed by afirst processor; and the conflicting access of the memory location is anintervening store to the memory location by a second processor.
 6. Themethod of claim 1, further comprising: logging, by hardware, informationabout the reinitiated hardware transaction including, at least, theinitiated hardware transaction attempt number and information about theconflicting access.
 7. The method of claim 1, further comprising:initiating a second hardware transaction in the computing environment,wherein the second hardware transaction accesses a second memorylocation; detecting a second conflicting access of the second memorylocation while executing the second hardware transaction; aborting thesecond hardware transaction based on the conflicting access of thesecond memory location; determining, by hardware, that the conflictingaccess of the second memory location is a non-transient condition; andresponsive to determining that the conflicting access of the memorylocation is a non-transient condition, aborting the second hardwaretransaction.